The Design &
Analysis of Algorithms
~~ I_ND EDIIIQN ~
Introduction to
*
The Design &
Analysis of Algorithms
IND EDITION "
Introduction to
~
~
~
Anany Levitin
Villanova University
Boston San Francisco New York
London Toronto Sydney Tokyo Singapore Madrid
Mexico City Munich Paris Cape Town Hong Kong Montreal
Publisher
Acquisitions Editor
Project Editor
Production Supervisor
Marketing Manager
Project Management & Composition
Technical Art
Copyeditor
Cover Design
Cover Illustration
Senior Pre press Supervisor
Greg Tobin
Matt Goldstein
Katherine Harutunian
Marilyn Lloyd
Michelle Brown
Windfall Software, using ZzT£X
George Nichols
Richard Camp
Joyce Cosentino Wells
Jennifer M. Kohnke
Caroline Fell
Access the latest information about Addison-Wesley titles from our World Wide Web site:
http://www.aw-bc.com/cornputing
Many of the designations used by manufacturers and sellers to distinguish their products
are claimed as trademarks. Where those designations appear in this book, and AddisonWesley was aware of a trademark claim, the designations have been printed in initial caps
or all caps.
If you purchased this book within the United States or Canada you should be aware that
it has been wrongfully imported without the approval of the Publisher or the Author.
Copyright© 2007 by Pearson Education, Inc.
All rights reserved. No part of this publication may be reproduced, stored in a retrieval
system, or transmitted, in any form or by any means, electronic, mechanical, photocopying,
recording, or otherwise, without the prior written permission of the publisher. Printed in
the United States of America.
ISBN 0-321-36413-9
3 4 5 6 7 8 9 10-DOH-09 08 07
Brief Contents
Preface
xvii
1 Introduction
1
41
97
123
157
197
249
279
307
335
379
415
2 Fundamentals of the Analysis of Algorithm Efficiency
3 Brute Force
4 Divide-and-Conquer
5 Decrease-and-Conquer
6 Transform-and-Conquer
1 Space and Time Tradeoffs
8 Dynamic Programming
9 Greedy Technique
1 0 Iterative Improvement
11 limitations of Algorithm Power
12 Coping with the Limitations of Algorithm Power
Epilogue
465
APPENDIX A
Useful Formulas for the Analysis of Algorithms
469
APPENDIX B
Short Tutorial on Recurrence Relations
473
Bibliography
487
497
541
Hints to Exercises
Index
v
r
.
.
I;
I!
Contents
Preface
xvii
1 Introduction
1
1.1 What is an Algorithm?
3
Exercises 1.1
1.2 Fundamentals of Algorithmic Problem Solving
8
Understanding the Problem
Ascertaining the Capabilities of a Computational Device
Choosing between Exact and Approximate Problem Solving
Deciding on Appropriate Data Structures
Algorithm Design Techniques
Methods of Specifying an Algorithm
Proving an Algorithm's Correctness
Analyzing an Algorithm
Coding an Algorithm
9
9
11
11
12
12
12
13
14
15
Exercises 1.2
17
1.3 Important Problem Types
Sorting
Searching
String Processing
Graph Problems
Combinatorial Problems
Geometric Problems
Numerical Problems
19
19
20
21
21
22
22
23
Exercises 1.3
23
vii
---~----·-·
I
viii
Contents
jli
;!'
rll
!:!
II:.,
il
,.'II
I!
1:1
1.4 Fundamental Data Structures
Linear Data Structures
Graphs
Trees
Sets and Dictionaries
Exercises 1.4
Summary
26
26
28
32
36
38
39
f:i
I'II:
1\,
I'-,,!
i!i
if!
li1
III
I
!
'
'
fundamentals of the Analysis of Algorithm
2 Efficiency
2.1 Analysis Framework
Measuring an Input's Size
Units for Measuring Running Time
Orders of Growth
Worst-Case. Best-Case, and Average-Case Efficiencies
Recapitulation of the Analysis Framework
Exercises 2.1
2.2 Asymptotic Notations and Basic Efficiency Classes
Informal Introduction
a-notation
0-notation
e-notation
Useful Property Involving the Asymptotic Notations
Using Limits for Comparing Orders of Growth
Basic Efficiency Classes
Exercises 2.2
41
42
43
44
45
47
50
50
52
52
53
54
55
56
57
58
59
2.3 Mathematical Analysis of Non recursive Algorithms
Exercises 2.3
61
67
2.4 Mathematical Analysis of Recursive Algorithms
Exercises 2.4
69
76
2.5 Example: Fibonacci Numbers
78
79
80
83
Explicit Formula for the nth Fibonacci Number
Algorithms for Computing Fibonacci Numbers
Exercises 2.5
Contents
2.6 Empirical Analysis of Algorithms
Exercises 2.6
2.7 Algorithm Visualization
ix
84
90
91
Summary
95
3 Brute Force
97
3.1 Selection Sort and Bubble Sort
Selection Sort
Bubble Sort
98
98
100
Exercises 3.1
3.2 Sequential Search and Brute-Force String Matching
Sequential Search
Brute-Force String Matching
Exercises 3.2
3.3 Closest-Pair and Convex-Hull Problems by Brute Force
Closest-Pair Problem
Convex-Hull Problem
Exercises 3.3
3.4 Exhaustive Search
Traveling Salesman Problem
Knapsack Problem
Assignment Problem
102
103
103
104
105
107
107
108
112
114
114
115
116
Exercises 3.4
119
Summary
120
4 Divide-and-Conquer
123
4.1 Mergesort
125
Exercises 4.1
128
4.2 Quicksort
129
Exercises 4.2
134
4.3 Binary Search
135
Exercises 4.3
138
r
111
)(
Contents
I
'
i::!
I:]
,,
4.4 Binary Tree Traversals and Related Properties
Exercises 4.4
139
142
4.5 Multiplication of Large Integers and Strassen's
Matrix Multiplication
Multiplication of Large Integers
Strassen's Matrix Multiplication
144
144
146
Exercises 4.5
l:i
i''
,.,
II
II
II
!I
I
4.6 Closest-Pair and Convex-Hull Problems by
Divide-and-Conquer
Closest-Pair Problem
Convex-Hull Problem
148
Exercises 4.6
149
149
150
154
Summary
155
:,f,,
•i
.,
Iii
iii
I!
5 Decrease-and-Conquer
157
160
:l'I;i
l!i
5.1 Insertion Sort
Exercises 5.1
163
'I
·:
5.2 Depth-First Search and Breadth-First Search
Depth-First Search
Breadth-First Search
Exercises 5.2
164
165
167
170
II
5.3 Topological Sorting
Exercises 5.3
172
176
5.4 Algorithms for Generating Combinatorial Objects
Generating Permutations
Generating Subsets
Exercises 5.4
177
178
180
181
5.5 Decrease-by-a-Constant-Factor Algorithms
Fake-Coin Problem
Multiplication a Ia Russe
Josephus Problem
183
183
184
185
187
',1
!I
!'I
j,
II
!I
il
,,!i
i!
,,
lr
f.!
lr'I
Exercises 5.5
5.6 Variable-Size-Decrease Algorithms
Computing a Median and the Selection Problem
188
188
Contents
xi
Interpolation Search
Searching and Insertion in a Binary Search Tree
The Game of Nim
190
191
192
Exercises 5.6
194
Summary
195
6 Transform-and-Conquer
197
6.1 Presorting
198
Exercises 6.1
201
6.2 Gaussian Elimination
Computing a Matrix Inverse
Computing a Determinant
203
208
209
210
Exercises 6.2
212
LU Decomposition and Other Applications
6.3 Balanced Search Trees
AVL Trees
2-3 Trees
214
214
218
Exercises 6.3
222
6.4 Heaps and Heapsort
Notion of the Heap
Heapsort
223
223
227
Exercises 6.4
229
6.5 Horner's Rule and Binary Exponentiation
Horner's Rule
Binary Exponentiation
230
231
233
Exercises 6.5
236
6.6 Problem Reduction
Computing the Least Common Multiple
Counting Paths in a Graph
Reduction of Optimization Problems
Linear Programming
Reduction to Graph Problerns
237
238
239
240
241
244
Exercises 6.6
245
Summary
247
8 Dynamic Programming
279
8.1 Computing a Binomial Coefficient
Exercises 8.1
280
282
8.2 Warshall's and Floyd's Algorithms
284
284
288
292
Warshall"s Algorithm
Floyd"s Algorithm for the All-Pairs Shortest-Paths Problem
Exercises 8.2
8.3 Optimal Binary Search Trees
Exercises 8.3
293
298
8.4 The Knapsack Problem and Memory Functions
299
Memory Functions
301
Exercises 8.4
Summary
303
304
Contents
Greedy Technique
xiii
307
9.1 Prim's Algorithm
308
Exercises 9.1
313
9.2 Kruskal's Algorithm
Disjoint Subsets and Union-Find Algorithms
Exercises 9.2
9.3 Dijkstra's Algorithm
315
317
321
323
Exercises 9.3
327
9.4 Huffman Trees
328
Exercises 9.4
332
Summary
333
10 Iterative Improvement
335
10.1 The Simplex Method
Geometric Interpretation of Linear Programming
An Outline of the Simplex Method
Further Notes on the Simplex Method
336
337
341
347
Exercises 10.1
10.2 The Maximum-Flow Problem
Exercises 10.2
10.3 Maximum Matching in Bipartite Graphs
Exercises 10.3
10.4 The Stable Marriage Problem
349
351
361
363
369
371
Exercises 10.4
375
Summary
376
11 Limitations of Algorithm Power
379
11.1 Lower-Bound Arguments
Trivial Lower Bounds
Information-Theoretic Arguments
380
381
382
~~
xiv
Contents
Adversary Arguments
Problem Reduction
382
383
Exercises 11.1
385
11.2 Decision Trees
Decision Trees for Sorting Algorithms
Decision Trees for Searching a Sorted Array
Exercises 11.2
11.3 P, NP, and NP-complete Problems
P and NP Problems
NP-Complete Problems
Exercises 11.3
11.4 Challenges of Numerical Algorithms
386
387
390
392
393
394
398
401
404
Exercises 11.4
411
Summary
413
12 Coping with the Limitations of Algorithm Power
12.1 Backtracking
n-Oueens Problem
Hamiltonian Circuit Problem
Subset-Sum Problem
General Remarks
Exercises 12.1
12.2 Branch-and-Bound
Assignment Problem
Knapsack Problem
Traveling Salesman Problem
Exercises 12.2
12.3 Approximation Algorithms for NP-hard Problems
Approximation Algorithms for the Traveling Salesman Problem
Approximation Algorithms for the Knapsack Problem
Exercises 12.3
12.4 Algorithms for Solving Nonlinear Equations
Bisection Method
Method of False Position
415
416
417
418
419
421
422
424
425
428
430
432
434
436
446
451
452
454
457
Contents
XV
Newton's Method
458
Exercises 12.4
461
Summary
462
Epilogue
465
APPENDIX A
Useful formulas for the Analysis of Algorithms
469
Properties of Logarithms
Combinatorics
Important Summation Formulas
Sum Manipulation Rules
Approximation of a Sum by a Definite Integral
Floor and Ceiling Formulas
Miscellaneous
469
469
470
470
471
471
471
APPENDIX 18
Short Tutorial on Recurrence Relations
473
Sequences and Recurrence Relations
Methods for Solving Recurrence Relations
Common Recurrence Types in Algorithm Analysis
473
474
479
Bibliography
487
Hints to Exercises
497
Index
541
r
I
Preface
The most valuable acquisitions in a scientific or technical education are the
general-purpose mental tools which remain serviceable for a life-time.
-George Forsythe, "What to do till the computer scientist comes," 1968
A
lgorithms play the central role in both the science and practice of computing.
Recognition of this fact has Jed to the appearance of a considerable number
of textbooks on the subject. By and large, they follow one of two alternatives
in presenting algorithms. One classifies algorithms according to a problem type.
Such a book would have separate chapters on algorithms for sorting, searching,
graphs, and so on. The advantage of this approach is that it allows an immediate
comparison of, say, the efficiency of different algorithms for the same problem.
The drawback of this approach is that it emphasizes problem types at the expense
of algorithm design techniques.
The second alternative organizes the presentation around algorithm design
techniques. In this organization, algorithms from different areas of computing are
grouped together if they have the same design approach. I share the belief of many
(e.g., [BaY95]) that this organization is more appropriate for a basic course on the
design and analysis of algorithms. There are three principal reasons for emphasis on algorithm design techniques. First, these techniques provide a student with
tools for designing algorithms for new problems. This makes learning algorithm
design techniques a very valuable endeavor from the practical standpoint. Second,
they seek to classify multitudes of known algorithms according to an underlying
design idea. Learning to see such commonality among algorithms from different
application areas should be a major goal of computer science education. After
all, every science considers classification of its principal subject as a major if not
the central point of its discipline. Third, in my opinion, algorithm design techniques have considerable utility as general problem solving strategies, applicable
to problems beyond computing.
xvii
L
xvm
Preface
Unfortunately, the traditional classification of algorithm design techniques
has several serious shortcomings from both theoretical and educational points
of view. The most significant of these shortcomings is the failure to classify many
important algorithms. This limitation has forced the authors of textbooks to depart
from the design technique organization and to include chapters dealing with
specific problem types. Such a switch leads to a loss of course coherence and almost
unavoidably creates confusion in students' minds.
New Taxonomy of Algorithm Design Techniques
My frustration with the shortcomings of the traditional classification of algorithm
design techniques has motivated me to develop a new taxonomy of them [Lev99],
which is the basis of this book. Here are the principal advantages of the new
taxonomy:
111
Ill
'"
'"
The new taxonomy is more comprehensive than the traditional one. It includes several strategies-brute force, decrease-and-conquer, transform-andconquer, space and time tradeoffs, and iterative improvement-that are rarely
if ever recognized as important design paradigms.
The new taxonomy covers naturally many classic algorithms (Euclid's algorithm, heapsort, search trees, hashing, topological sorting, Gaussian elimination, Horner's rule, to name a few) that the traditional taxonomy cannot
classify. As a result, the new taxonomy makes it possible to present the standard body of classic algorithms in a unified and coherent fashion.
It naturally accommodates the existence of important varieties of several
design techniques. (For example, it recognizes three variations of decreaseand-conquer and three variations of transform-and-conquer.)
It is better aligned with analytical methods for efficiency analysis (see Appendix B).
Design Techniques as General Problem Solving Strategies
Most applications of the design techniques in the book are to classic problems of
computer science. (The only innovation here is the inclusion of some material on
numerical algorithms, which are covered within the same general framework.)
But these design techniques can be considered general problem solving tools,
whose applications are not limited to traditional computing and mathematical
problems. Two factors make this point particularly important. First, more and
more computing applications go beyond the traditional domain, and there are
reasons to believe that this trend will strengthen in the future. Second, developing
students' problem solving skills has come to be recognized as a major goal of
college education. Among all the courses in a computer science curriculum, a
Preface
xix
course on the design and analysis of algorithms is uniquely suitable for this task
because it can offer a student specific strategies for solving problems.
I am not proposing that a course on the design and analysis of algorithms
should become a course on general problem solving. But I do believe that the
unique opportunity provided by studying the design and analysis of algorithms
should not be missed. Toward this goal, the book includes applications to puzzles
and puzzle-like games. Although using puzzles in teaching algorithms is certainly
not a new idea, the book tries to do so systematically by going well beyond a few
standard examples.
Textbook Pedagogy
My goal was to write a text that would not trivialize the subject but would still be
readable by most students on their own. Here are some of the things done toward
this objective.
"
"
"
"
"
Sharing the opinion of George Forsythe expressed in the epigraph, I have
sought to stress major ideas underlying the design and analysis of algorithms.
In choosing specific algorithms to illustrate these ideas, I limited the number
of covered algorithms to those that most clearly demonstrate an underlying
design technique or analysis method. Fortunately, most classic algorithms
satisfy this criterion.
In Chapter 2, which is devoted to efficiency analysis, the methods used for
analyzing nonrecursive algorithms are separated from those typically used for
analyzing recursive algorithms. The chapter also includes sections devoted to
empirical analysis and algorithm visualization.
The narrative is systematically interrupted by questions to the reader. Some
of them are asked rhetorically, in anticipation of a concern or doubt, and are
answered immediately. The goal of the others is to prevent the reader from
drifting through the text without a satisfactory level of comprehension.
Each chapter ends with a summary recapping the most important concepts
and results discussed in the chapter.
The book contains about 700 exercises. Some of them are drills; others make
important points about the material covered in the body of the text or introduce algorithms not covered there at all. A few exercises take advantage of
Internet resources. More difficult problems-there are not many of them-are
marked with a special symbol in the Instructor's Manual. (Because designating problems as difficult may discourage some students from trying to tackle
them, problems are not marked in the book itself.) Puzzles, games, and puzzlelike questions are marked in the exercises with a special icon.
The book provides hints to all the exercises. Detailed solutions, except for
programming projects, are provided in the Instructor's Manual, available
r
XX
Preface
to qualified adopters through Addison-Wesley's Instructor Resource Center. (Please contact your local Addison-Wesley sales representative or visit
www.aw.com/irc to access this material.) Slides in PowerPoint are available to
all readers of this book at www.aw.comfcssupport.
Changes for the Second Edition
The most important change in the second edition is a new chapter (Chapter 10)
dealing with the iterative-improvement technique. There are several reasons for
this addition.
10
10
'"
'"
Some of the most important algorithms in computer science, both from a
theoretical and practical standpoint, are based on this technique.
Iterative improvement is the principal alternative to the greedy strategy already covered in the book.
Since linear programming is discussed in Chapter 6, it is natural to cover its
principal algorithm (the simplex method) as well.
Iterative improvement is the only design technique of the new taxonomy that
was not included in the first edition. Therefore, its inclusion could also be
justified by the sheer need for completeness.
The four sections of the new chapter cover the following topics: the simplex
method, maximum network flows, maximum cardinality matching in bipartite
graphs, and the stable marriage problem. Each of these sections can be studied
independently of the rest, as is the case for the other chapters of the book.
There are two other additions in this edition worth mentioning. The first is
a new subsection about the game of Nim, included in Section 5.6 on variabledecrease algorithms. The second is an expanded coverage of approximation algorithms for the traveling salesman problem to give a better account of the remarkable progress achieved in this area.
I also added about 80 new problems to the exercises. (A list of the new
problems is available on the book's Web site at www.aw-bc.com/info/levitin.)
About half of them are added to the old chapters, and a majority are puzzles or
puzzle-like questions. Judging by the feedback received, the inclusion of puzzles
and games in tbe exercises has been welcome by the book's readers, including
instructors and students using it as a textbook. As in the first edition, only the
puzzles related to the discussed topics were included in the exercises.
Prerequisites
The book assumes that a reader has gone through an introductory programming
course and a standard course on discrete structures. With such a background,
he or she should be able to handle the book's material without undue difficulty.
Still, fundamental data structures, necessary summation formulas, and recurrence
relations are reviewed in Section 1.4, Appendix A, and Appendix B, respectively.
Preface
xxi
Calculus is used in only three sections (Sections 2.2, 11.4, and 12.4) and to a very
limited degree; if students lack calculus as an assured part of their background,
the portions of these three sections that involve calculus can be omitted without
hindering their understanding of the rest of the material.
Use in the Curriculum
The book can serve as a textbook for a basic course on design and analysis
of algorithms organized around algorithm design techniques. It might contain
too much material for a typical one-semester course. By and large, portions of
Chapters 3 through 12 can be skipped without the danger of making later parts
of the book incomprehensible to the reader. Any portion of the book can be
assigned for self-study. In particular, Sections 2.6 and 2.7 on empirical analysis and
algorithm visualization, respectively, can be assigned in conjunction with projects.
Here is a possible plan for a one-semester course; it assumes a 40-class meeting
format.
Lecture
Topic
Sections
1
2,3
4
5,6
7
8
9-10
Introduction
1.1-1.3
2.1,2.2
2.3
2.4, 2.5 (+ App. B)
3.1, 3.2 (+ 3.3)
3.4
4.1-4.3
11
12-14
15
16
17-19
20
21
22-24
25-27
28-30
Analysis framework; 0, 8, n notations
Mathematical analysis of nonrecursive algorithms
Mathematical analysis of recursive algorithms
Brute-force algorithms
Exhaustive search
Divide-and-conquer: mergesort, quicksort, binary
search
Other divide-and-conquer examples
Decrease-by-one: insertion sort, DFS and BFS,
topological sorting
Decrease-by-a-constant-factor algorithms
Variable-size-decrease algorithms
Instance simplification: presorting, Gaussian
elimination, balanced search trees
Representation change: heaps and heapsort or
Horner's rule and binary exponentiation
Problem reduction
Space-time tradeoffs: string matching, hashing,
B-trees
Dynamic programming algorithms
Greedy algorithms: Prim's, Kruskal's, Dijkstra's,
Huffman's
4.4 or 4.5 or 4.6
5.1-5.3
5.5
5.6
6.1-6.3
6.4 or
6.5
6.6
7.2-7.4
3 from 8.1-8.4
9.1-9.4
xxii
Preface
Lecture
Topic
Sections
31-33
34
35
36
37
38
39
40
Iterative improvement algorithms
Lower-bound arguments
Decision trees
P, NP, and NP-complete problems
Numerical algorithms
3 from 10.1-10.4
11.1
11.2
11.3
11.4 ( + 12.4)
12.1
12.2
12.3
Backtracking
Branch-and-bound
Approximation algorithms for NP-hard problems
Acknowledgments
I would like to express my gratitude to the many people who have shared with
me their own and their students' opinions about the first edition and suggested
improvements and corrections. I have especially benefitted from the comments
by Eric Bach (University of Wisconsin), Stephen Bique (University of AlaskaFairbanks), Walter Hower (Albstadt-Sigmaringen University), Dan Joyce (Villanova University), Steven Lindell (Haverford College), and Mary-Angela Papalaskari (Villanova University).
My thanks also go to Edward L. Bosworth (Columbus State University),
Philip Dorin (Loyola Marymount University), William Ford (University of the
Pacific), Jonathan Goldstine (Penn State University), Chin-Cheng Hung (Southern Polytechnic State University), and Brian Rosmaita (Hamilton College )-the
reviewers of the chapter on iterative-improvement algorithms-for supporting the
inclusion of this chapter in the new edition and for their ideas about presenting
these important algorithms.
In addition, I would like to convey my appreciation to the following reviewers: David Avis (McGill University), Gordon Beavers (University of Arkansas),
Dinesh Mehta (Colorado School of Mines), Feridoon Moinian (Cameron University), Stephan Olariu (Old Dominion University), Stanley Selkow (WPI), Jerry
Trahan (Louisiana State University), and Yang Wang (Missouri State University).
I gratefully acknowledge the help I have received from my former student,
Rakhi Gandhi, whose assistance has certainly decreased the number of typos and
errors in the book.
I appreciate the contributions of all the reviewers of the first edition: Simon Berkovich (George Washington University), Richard Borie (University of
Alabama), Douglas M. Campbell (Brigham Young University), Bin Cong (California State University, Fullerton), Steve Homer (Boston University), Roland
Htibscher (Auburn University), Sukhamay Kundu (Louisiana State University),
Sheau-Dong Lang (University of Central Florida), John C. Lusth (University of
Arkansas), John F. Meyer (University of Michigan), Steven R. Seidel (Michigan
Preface
xxiii
Technological University), Ali Shokoufandeh (Drexel University), and George
H. Williams (Union College).
My thanks go to all the people at Addison-Wesley and their associates who
worked on my book. I am especially grateful to my editor Matt Goldstein, Michelle
Brown, Katherine Harutunian, Marilyn Lloyd, Joyce Wells, and Paul Anagnostopoulos of Windfall Software. I was lucky to have Richard Camp as the copy editor;
his numerous suggestions have definitely improved the book.
Finally, I am indebted to two members of my family. Living with a spouse
writing a book is probably more trying than doing the actual writing. My wife,
Maria, lived through several years of this, helping me any way she could. And help
she did: over 400 figures in the book and the Instructor's Manual were created by
her. My daughter Miriam has been my English prose guru over many years. She
not only read large portions of the book, but she was also instrumental in finding
the chapter epigraphs.
Anany Levitin
anany.levitin@villanova.edu
December 2005
The Design &
Analysis of Algorithms
fiND EDFFIC>N .
Introduction to
i
,I
:·!
:iJ
II.I
iil!
:i!
I'
il
[i
1
Introduction
Two ideas lie gleaming on the jeweler's velvet. The first is the calculus,
the second, the algorithm. The calculus and the rich body of mathematical
analysis to which it gave rise made modern science possible; but it has been
the algorithm that has rnade possible the modern world.
-David Berlinski, The Advent of the Algorithm, 2000
W
hy do you need to study algorithms? If you are going to be a computer
professional, there are both practical and theoretical reasons to study algorithms. From a practical standpoint, you have to know a standard set of important
algorithms from different areas of computing; in addition, you should be able to
design new algorithms and analyze their efficiency. From the theoretical standpoint, the study of algorithms, sometimes called algorithmics, has come to be
recognized as the cornerstone of computer science. David Hare!, in his delightful
book pointedly titled Algorithmics: the Spirit of Computing, put it as follows:
Algorithmics is more than a branch of computer science. It is the core of
computer science, and, in all fairness, can be said to be relevant to most of
science, business, and technology. [Har92], p. 6.
But even if you are not a student in a computing-related program, there are
compelling reasons to study algorithms. To put it bluntly, computer programs
would not exist without algorithms. And with computer applications becoming
indispensable in almost all aspects of our professional and personal lives, studying
algorithms becomes a necessity for more and more people.
Another reason for studying algorithms is their usefulness in developing analytical skills. After all, algorithms can be seen as special kinds of solutions to
problems-not answers but precisely defined procedures for getting answers. Consequently, specific algorithm design techniques can be interpreted as problemsolving strategies that can be useful regardless of whether a computer is involved.
Of course, the precision inherently imposed by algorithmic thinking limits the
kinds of problems that can be solved with an algorithm. You will not find, for
example, an algorithm for living a happy life or becoming rich and famous. On
1
.
w
.
'
2
Introduction
i!
I'
I
A person well-trained in computer science knows how to deal with algorithms:
how to construct them, manipulate them, understand them, analyze them.
This knowledge is preparation for much more than writing good computer
programs; it is a general-purpose mental tool that will be a definite aid to
the understanding of other subjects, whether they be chemistry, linguistics,
or music, etc. The reason for this may be understood in the following way:
It has often been said that a person does not really understand something
until after teaching it to someone else. Actually, a person does not really
understand something until after teaching it to a computer, i.e., expressing
it as an algorithm ... An attempt to formalize things as algorithms leads to
a much deeper understanding than if we simply try to comprehend things in
the traditional way. [Knu96], p. 9.
I
1:
I
IIi
f1
'I
I
II
'I
~i
·i,'
'I
''!
!i
i'
'10
!J
1
the other hand, this required precision has an important educational advantage.
Donald Knuth, one of the most prominent computer scientists in the history of
algorithmics, put it as follows:
We take up the notion of algorithm in Section 1.1. As examples, we use three
algorithms for the same problem: computing the greatest common divisor. There
are several reasons for this choice. First, it deals with a problem familiar to everybody from their middle-school days. Second, it makes the important point that
the same problem can often be solved by several algorithms. Quite typically, these
algorithms differ in their idea, level of sophistication, and efficiency. Third, one of
these algorithms deserves to be introduced first, both because of its age~it appeared in Euclid's famous treatise more than two thousand years ago~and its
enduring power and importance. Finally, the middle-school procedure for computing the greatest common divisor allows us to highlight a critical requirement
every algorithm must satisfy.
Section 1.2 deals with algorithmic problem solving. There we discuss several
important issues related to the design and analysis of algorithms. The different
aspects of algorithmic problem solving range from analysis of the problem and the
means of expressing an algorithm to establishing its correctness and analyzing its
efficiency. The section does not contain a magic recipe for designing an algorithm
for an arbitrary problem. It is a well-established fact that such a recipe does not
exist. Still, the material of Section 1.2 should be useful for organizing your work
on designing and analyzing algorithms.
Section 1.3 is devoted to a few problem types that have proven to be particularly important to the study of algorithms and their application. In fact, there
are textbooks (e.g., [Sed88]) organized around such problem types. I hold the
view~shared by many others~that an organization based on algorithm design
techniques is superior. In any case, it is very important to be aware of the principal problem types. Not only are they the most commonly encountered problem
types in real-life applications, they are used throughout the book to demonstrate
particular algorithm design techniques.
I
t
1.1
What is an Algorithm?
3
Section 1.4 contains a review of fundamental data structures. It is meant to
serve as a reference rather than a deliberate discussion of this topic. If you need
a more detailed exposition, there is a wealth of good books on the subject, most
of them tailored to a particular programming language.
1.1
What is an Algorithm?
Although there is no universally agreed-on wording to describe this notion, there
is general agreement about what the concept means:
An algorithm is a sequence of unambiguous instructions for solving a
problem, i.e., for obtaining a required output for any legitimate input in
a finite amount of time.
This definition can be illustrated by a simple diagram (Figure 1.1 ).
The reference to "instructions" in the definition implies that there is something or someone capable of understanding and following the instructions given.
We call this a "computer," keeping in mind that before the electronic computer
was invented, the word "computer" meant a human being involved in performing numeric calculations. Nowadays, of course, "computers" are those ubiquitous
electronic devices that have become indispensable in almost everything we do.
Note, however, that although the majority of algorithms are indeed intended for
eventual computer implementation, the notion of algorithm does not depend on
such an assumption.
As examples illustrating the notion of algorithm, we consider in this section
three methods for solving the same problem: computing the greatest common
divisor of two integers. These examples will help us to illustrate several important
points:
"'
The nonambiguity requirement for each step of an algorithm cannot be com·
promised.
problem
1
algorithm
input
---.[==·~c~om~p~ut~e~r·~·=}----+ output
FIGURE 1.1 Notion of algorithm
4
Introduction
"
,.
"
II!
The range of inputs for which an algorithm works has to be specified carefully.
The same algorithm can be represented in several different ways.
Several algorithms for solving the same problem may exist.
Algorithms for the same problem can be based on very different ideas and
can solve the problem with dramatically different speeds.
Recall that the greatest common divisor of two nonnegative, not -both-zero
integers m and n, denoted gcd(m, n), is defined as the largest integer that divides
both m and n evenly, i.e., with a remainder of zero. Euclid of Alexandria (third
century B.C.) outlined an algorithm for solving this problem in one of the volumes
of his Elements, most famous for its systematic exposition of geometry. In modern
terms, Euclid's algorithm is based on applying repeatedly the equality
gcd(m, n) = gcd(n, m mod n)
(where m mod n is the remainder of the division of m by n) until m mod n is equal
to 0; since gcd(m, 0) = m (why?), the last value of m is also the greatest common
divisor of the initial m and n.
For example, gcd(60, 24) can be computed as follows:
gcd(60, 24) = gcd(24, 12) = gcd(12, 0) = 12.
(If you are not impressed by this algorithm, try finding the greatest common divisor
of larger numbers such as those in Problem 5 of Exercises 1.1.)
Here is a more structured description of this algorithm:
Euclid's algorithm for computing gcd(m, n)
Step 1 If n = 0, return the value of m as the answer and stop; otherwise,
proceed to Step 2.
Step 2 Divide m by n and assign the value of the remainder to r.
Step 3 Assign the value of n tom and the value of r ton. Go to Step 1.
Alternatively, we can express the same algorithm in a pseudocode:
ALGORITHM
Euclid(m, n)
//Computes gcd(m, n) by Euclid's algorithm
!!Input: Two nonnegative, not -both-zero integers m and n
//Output: Greatest common divisor of m and n
while n oJ 0 do
r +-m modn
m +-n
n +---- r
return m
How do we know that Euclid's algorithm eventually comes to a stop? This
follows from the observation that the second number of the pair gets smaller with
1.1
What is an Algorithm?
5
each iteration and it cannot become negative. Indeed, the new value of n on the
next iteration ism mod n, which is always smaller than n. Hence, the value of the
second number in the pair eventually becomes 0, and the algorithm stops.
Just as with many other problems, there are several algorithms for computing
the greatest common divisor. Let us look at the other two methods for this problem. The first is simply based on the definition of the greatest common divisor of
m and n as the largest integer that divides both numbers evenly. Obviously, such
a common divisor cannot be greater than the smaller of these numbers, which we
will denote by t = min{m, n ). So we can start by checking whether t divides both
m and n: if it does, t is the answer; if it does not, we simply decrease t by 1 and
try again. (How do we know that the process will eventually stop?) For example,
for numbers 60 and 24, the algorithm will try first 24, then 23, and so on until it
reaches 12, where it stops.
Consecutive integer checking algorithm for computing gcd(m, n)
Step 1 Assign the value of min{m, n) tot.
Step 2 Divide m by t. If the remainder of this division is 0, go to Step 3;
otherwise, go to Step 4.
Step 3 Divide n by t. If the remainder of this division is 0, return the value of
t as the answer and stop; otherwise, proceed to Step 4.
Step 4 Decrease the value oft by 1. Go to Step 2.
Note that unlike Euclid's algorithm, this algorithm, in the form presented,
does not work correctly when one of its input numbers is zero. This example
illustrates why it is so important to specify the range of an algorithm's inputs
explicitly and carefully.
The third procedure for finding the greatest common divisor should be familiar to you from middle school.
Middle-school procedure for computing gcd(m, n)
Step 1 Find the prime factors of m.
Step 2 Find the prime factors of n.
Step 3 Identify all the common factors in the two prime expansions found in
Step 1 and Step 2. (If p is a common factor occurring Pm and Pn times
in m and n, respectively, it should be repeated min{pm, p,) times.)
Step 4 Compute the product of all the common factors and return it as the
greatest common divisor of the numbers given.
Thus, for the numbers 60 and 24, we get
60=2·2·3·5
24=2·2·2·3
gcd(60, 24) = 2. 2. 3 = 12.
'7
6
Introduction
Nostalgia for the days when we learned this method should not prevent us
from noting that the last procedure is much more complex and slower than Euclid's
algorithm. (We will discuss methods for finding and comparing running times of
algorithms in the next chapter.) In addition to inferior efficiency, the middle-school
procedure does not qualify, in the form presented, as a legitimate algorithm. Why?
Because the prime factorization steps are not defined unambiguously: they require
a list of prime numbers, and I strongly suspect that your middle-school teacher
did not explain how to obtain such a list. You undoubtedly agree that this is not
a matter of unnecessary nitpicking. Unless this issue is resolved, we cannot, say,
write a program implementing this procedure. (Incidentally, Step 3 is also not
defined clearly enough. Its ambiguity is much easier to rectify than that of the
factorization steps, however. How would you find common elements in two sorted
lists?)
So let us introduce a simple algorithm for generating consecutive primes
not exceeding any given integer n. It was probably invented in ancient Greece
and is known as the sieve of Eratosthenes (ca. 200 B.C.). The algorithm starts by
initializing a list of prime candidates with consecutive integers from 2 ton. Then,
on the first iteration of the algorithm, it eliminates from the list all multiples of
2, i.e., 4, 6, and so on. Then it moves to the next item on the list, which is 3,
and eliminates its multiples. (In this straightforward version, there is an overhead
because some numbers, such as 6, are eliminated more than once.) No pass for
number 4 is needed: since 4 itself and all its multiples are also multiples of 2, they
were already eliminated on a previous pass. (By similar reasoning, we need not
consider multiples of any eliminated number.) The next remaining number on the
list, which is used on the third pass, is 5. The algorithm continues in this fashion
until no more numbers can be eliminated from the list. The remaining integers of
the list are the primes needed.
As an example, consider the application of the algorithm to finding the list of
primes not exceeding n = 25:
2 3 4 5 6 7 8 9 10 11 12 13 14 15 16 17 18 19 20 21 22 23 24 25
17
19
5
7
9
11
13
15
21
23
25
2 3
5
7
11
17
19
23
25
13
2 3
5
7
11
17
19
23
2 3
13
For this example, no more passes are needed because they would eliminate numbers already eliminated on previous iterations of the algorithm. The remaining
numbers on the list are the consecutive primes less than or equal to 25.
In general, what is the largest number p whose multiples can still remain on
the list? Before we answer this question, let us first note that if p is a number
whose multiples are being eliminated on the current pass, then the first multiple
we should consider is p · p because all its smaller multiples 2p, ... , (p- l)p have
been eliminated on earlier passes through the list. This observation helps to avoid
1.1
What is an Algorithm?
7
eliminating the same number more than once. Obviously, p · p should not be
greater than n, and therefore p cannot exceed ,jn rounded down (denoted ,jn
using the so-called floor function). We assume in the following pseudocode that
there is a function available for computing ,jn alternatively, we could check
the inequality p · p S n as the loop continuation condition there.
l J
l J;
ALGORITHM
Sieve(n)
//Implements the sieve of Eratosthenes
//Input: An integer n ": 2
//Output: Array L of all prime numbers less than or equal ton
for p-<-- 2 ton do A[p]-<- p
for p -<-- 2 to ,Jnj do //see note before pseudocode
if A[p] i' 0 II p hasn't been eliminated on previous passes
l
j-<-P*P
while j S n do
A[j] -<-- 0 //mark element as eliminated
j-<-j+p
//copy the remaining elements of A to array L of the primes
i-<--0
forp-<-2tondo
ifA[p]i'O
L[i]+- A[p]
i-<-i+1
return L
So now we can incorporate the sieve of Eratosthenes into the middle-school
procedure to get a legitimate algorithm for computing the greatest common divisor of two positive integers. Note that special care needs to be exercised if one
or both input numbers are equal to 1: because mathematicians do not consider
1 to be a prime number, strictly speaking, the method does not work for such
inputs.
Before we leave this section, one more comment is in order. The examples
considered in this section notwithstanding, the majority of algorithms in use
today-even those that are implemented as computer programs-do not deal
with mathematical problems. Look around for algorithms helping us through our
daily routines, both professional and personal. May this ubiquity of algorithms in
today's world strengthen your resolve to learn more about these fascinating
engines of the information age.
8
Introduction
-----Exercises 1.1
1. Do some research on al-Khorezmi (also al-Khwarizmi), the man from whose
name the word "algorithm" is derived. 1n particular, you should learn what
the origins of the words "algorithm)' and "algebra" have in common.
2. Given that the official purpose of the U.S. patent system is the promotion
of the "useful arts," do you think algorithms are patentable in this country?
Should they be?
3. a. Write down driving directions for going from your school to your home
with the precision required by an algorithm.
b. Write down a recipe for cooking your favorite dish with the precision
required by an algorithm.
:I
4. Design an algorithm for computing LftJ for any positive integer n. Besides
assignment and comparison, your algorithm may only use the four basic
arithmetical operations.
5. a. Find gcd(31415, 14142) by applying Euclid's algorithm.
b. Estimate how many times faster it will be to find gcd(31415, 14142) by
Euclid's algorithm compared with the algorithm based on checking consecutive integers from min{m, n) down to gcd(m, n).
6. Prove the equality gcd(m, n) = gcd(n, m mod n) for every pair of positive
integers m and n.
7. What does Euclid's algorithm do for a pair of numbers in which the first
number is smaller than the second one? What is the largest number of times
this can happen during the algorithm's execution on such an input?
8. a. What is the smallest number of divisions made by Euclid's algorithm
among all inputs 1 ::= m, n ::= 10?
b. What is the largest number of divisions made by Euclid's algorithm among
all inputs 1 ::= m, n ::= 10?
9. a. Euclid's algorithm, as presented in Euclid's treatise, uses subtractions
rather than integer divisions. Write a pseudocode for this version of Euclid's algorithm.
b. Euclid's game (see [Bog]) starts with two unequal positive numbers on the
board. Two players move in turn. On each move, a player has to write on
the board a positive number equal to the difference oftwo numbers already
on the board; this number must be new, i.e., different from all the numbers
already on the board. The player who cannot move loses the game. Should
you choose to move first or second in this game?
10. The extended Euclid's algorithm determines not only the greatest common
divisor d of two positive integers m and n but also integers (not necessarily
positive) x andy, such that mx + ny =d.
1.2
Fundamentals of Algorithmic Problem Solving
9
a. Look up a description of the extended Euclid's algorithm (see, e.g., [Knui],
p. 13) and implement it in the language of your choice.
b. Modify your program for finding integer solutions to the Diophantine
equation ax +by= c with any set of integer coefficients a, b, and c.
11. Locker doors There are n lockers in a hallway, numbered sequentially from
1 to n. Initially all the locker doors are closed. You make 11 passes by the
lockers, each time starting with locker #1. On the ith pass, i = 1, 2, ... , 11, you
toggle the door of every ith locker: if the door is closed, you open it; if it is
open, you close it. For example, after the first pass every door is open; on the
second pass you only toggle the even-numbered lockers (#2, #4, ... ) so that
after the second pass the even doors are closed and the odd ones are open;
the third time through, you close the door of locker #3 (opened from the first
pass), open the door of locker #6 (closed from the second pass), and so on.
After the last pass, which locker doors are open and which are closed? How
many of them are open?
1.2 Fundamentals of Algorithmic Problem Solving
Let us start by reiterating an important point made in the introduction to this
chapter:
We can consider algorithms to be procedural solutions to problems.
These solutions are not answers but specific instructions for getting answers. It is
this emphasis on precisely defmed constructive procedures that makes computer
science distinct from other disciplines. In particular, this distinguishes it from theoretical mathematics whose practitioners are typically satisfied with just proving
the existence of a solution to a problem and, possibly, investigating the solution's
properties.
We now list and briefly discuss a sequence of steps one typically goes through
in designing and analyzing an algorithm (Figure 1.2).
Understanding the Problem
From a practical perspective, the frrst thing you need to do before designing an
algorithm is to understand completely the problem given. Read the problem's
description carefully and ask questions if you have any doubts about the problem,
do a few small examples by hand, think about special cases, and ask questions
again if needed.
There are a few types of problems that arise in computing applications quite
often. We review them in the next section. If the problem in question is one of
them, you might be able to use a known algorithm for solving it. Of course, it
helps to understand how such an algorithm works and to know its strengths and
r
10
Introduction
Understand the problem
Decide on:
computational means,
exact vs. approximate solving,
data structure(s),
algorithm design technique
Design an algorithm
Prove correctness
Analyze the algorithm
Code the algorithm
FIGURE 1.2 Algorithm design and analysis process
weaknesses, especially if you have to choose among several available algorithms.
But often, you will not find a readily available algorithm and will have to design
your own. The sequence of steps outlined in this section should help you in this
exciting but not always easy task.
.
An input to an algorithm specifies an ';nstanc~ of the problem the algorithm
solves. It is very important to specify exactly the range of instances the algorithm
needs to handle. (As an example, recall the variations in the range of instances for
the three greatest common divisor algorithms discussed in the previous section.)
If you fail to do this, your algorithm may work correctly for a majority of inputs
but crash on some "boundary" value.,t-gemember that a correct algorithm is not
one that works most of the time, but
that works correctly for alllegitiJ!!aty
inputs.
Do not skimp on this first step of the algorithmic problem-solving process; if
you do, you will run the risk of unnecessary rework.
'One
1.2
Fundamentals of Algorithmic Problem Solving
11
Ascertaining the Capabilities of a Computational Device
Once you completely understand a problem. you need to ascertain the capabilities
of the computational device the algorithm is intended for. The vast majority of
algorithms in use today are still destined to be programmed for a computer closely
resembling the von Neumann machine-a computer architecture outlined by
the prominent Hungarian-American mathematician John von Neumann (19031957), in collaboration with A. Burks and H. Goldstine, in 1946. The essence of
this architecture is captured by the so-called random-access machine (RAM).
Its central assumption is that instructions are executed one after another, one
operation at a time. Accordingly, algorithms designed to be executed on such
machines are called sequential algorithms.
The central assumption of the RAM model does not hold for some newer
computers that can execute operations concurrently, i.e., in parallel. Algorithms
that take advantage of this capability are called parallel algorithms. Still, studying
the classic techniques for design and analysis of algorithms under the RAM model
remains the cornerstone of algorithmics for the foreseeable future.
Should you worry about the speed and amount of memory of a computer
at your disposal? If you are designing an algorithm as a scientific exercise, the
answer is a qualified no. As you will see in Section 2.1, most computer scientists
prefer to study algorithms in terms independent of specification parameters for
a particular computer. If you are designing an algorithm as a practical tool, the
answer may depend on a problem you need to solve. Even "slow" computers of
today are almost unimaginably fast. Consequently, in many situations, you need
not worry about a computer being too slow for the task. There are important
problems, however, that are very complex by their nature, have to process huge
volumes of data, or deal with applications where time is critical. In such situations,
it is imperative to be aware of the speed and memory available on a particular
computer system.
Choosing between Exact and Approximate Problem Solving
The next principal decision is to choose between solving the problem exactly or
solving it approximately. In the former case, an algorithm is called an exact algorithm; in the latter case, an algorithm is called an approximation algorithm. Why
would one opt for an approximation algorithm? First, there are important problems that simply cannot be solved exactly for most of their instances; examples
include extracting square roots, solving nonlinear equations, and evaluating definite integrals. Second, available algorithms for solving a problem exactly can be
unacceptably slow because of the problem's intrinsic complexity. Tbis happens, in
particular, for many problems involving a very large number of choices; you will
see examples of such difficult problems in Chapters 3, 11, and 12. Third, an approximation algorithm can be a part of a more sophisticated algorithm that solves
a problem exactly.
12
Introduction
Deciding on Appropriate Data Structures
Some algorithms do not demand any ingenuity in representing their inputs. But
others are, in fact, predicated on ingenious data structures. In addition, some
of the algorithm design techniques we shall discuss in Chapters 6 and 7 depend
intimately on structuring or restructuring data specifying a problem's instance.
Many years ago, an influential textbook proclaimed the fundamental importance
of both algorithms and data structures for computer programming by its very title:
Algorithms + Data Structures ~ Programs [Wir76]. In the new world of objectoriented programming, data structures remain crucially important for both design
and analysis of algorithms. We review basic data structures in Section 1.4.
Algorithm Design Techniques
Now, with all the components of the algorithmic problem solving in place, how do
you design an algorithm to solve a given problem? This is the main question this
book seeks to answer by teaching you several general design techniques.
What is an algorithm design technique?
An algorithm design technique (or "strategy" or "paradigm") is a general
approach to solving problems algorithmically that is applicable to a variety
of problems from different areas of computing.
Check this book's table of contents and you will see that a majority of its
chapters are devoted to individual design techniques. They distill a few key ideas
that have proven to be useful in designing algorithms. Learning these techniques
is of utmost importance for the following reasons.
First, they provide guidance for designing algorithms for new problems, i.e.,
problems for which there is no known satisfactory algorithm. Therefore-to use
the language of a famous proverb-learning such techniques is akin to learning
to fish as opposed to being given a fish caught by somebody else. It is not true, of
course, that each of these general techniques will be necessarily applicable to every
problem you may encounter. But taken together, they do constitute a powerful
collection of tools that you will find quite handy in your studies and work.
Second, algorithms are the cornerstone of computer science. Every science is
interested in classifying its principal subject, and computer science is no exception.
Algorithm design techniques make it possible to classify algorithms according
to an underlying design idea; therefore, they can serve as a natural way to both
categorize and study algorithms.
Methods of Specifying an Algorithm
Once you have designed an algorithm, you need to specify it in some fashion. In
Section 1.1, to give you an example, we described Euclid's algorithm in words (in
a free and also a step-by-step form) and in pseudocode. These are the two options
that are most widely used nowadays for specifying algorithms.
l
l
1.2 Fundamentals of Algorithmic Problem Solving
13
Using a natural language has an obvious appeal; however, the inherent ambiguity of any natural language makes a succinct and clear description of algorithms
surprisingly difficult. Nevertheless, being able to do this is an important skill that
you should strive to develop in the process of learning algorithms.
A pseudocode is a mixture of a natural language and programming languagelike constructs. A pseudocode is usually more precise than a natural language, and
its usage often yields more succinct algorithm descriptions. Surprisingly, computer
scientists have never agreed on a single form of pseudocode, leaving textbook
authors to design their own "dialects." Fortunately, these dialects are so close to
each other that anyone familiar with a modern programming language should be
able to understand them all.
This book's dialect was selected to cause minimal difficulty for a reader. For
the sake of simplicity, we omit declarations of variables and use indentation to
show the scope of such statements as for, if, and while. As you saw in the previous
section, we use an arrow <-- for the assignment operation and two slashes I I for
comments.
In the earlier days of computing, the dominant vehicle for specifying algorithms was a flowchart, a method of expressing an algorithm by a collection of
connected geometric shapes containing descriptions of the algorithm's steps. This
representation technique has proved to be inconvenient for all but very simple
algorithms; nowadays, it can be found only in old algorithm books.
The state of the art of computing has not yet reached a point where an
algorithm's description-whether in a natural language or a pseudocode-can be
fed into an electronic computer directly. Instead, it needs to be converted into a
computer program written in a particular computer language. We can look at such
a program as yet another way of specifying the algorithm, although it is preferable
to consider it as the algorithm's implementation.
Proving an Algorithm's Correctness
Once an algorithm has been specified, you have to prove its correctness. That is,
you have to prove that the algorithm yields a required result for every legitimate
input in a finite amount of time. For example, correctness of Euclid's algorithm
for computing the greatest common divisor stems from correctness of the equality
gcd(m, n) = gcd(n, m mod n) (which, in turn, needs a proof; see Problem 6 in
Exercises 1.1), the simple observation that the second number gets smaller on
every iteration of the algorithm, and the fact that the algorithm stops when the
second number becomes 0.
For some algorithms, a proof of correctness is quite easy; for others, it can
be quite complex. A common technique for proving correctness is to use mathematical induction because an algorithm's iterations provide a natural sequence of
steps needed for such proofs. It might be worth mentioning that although tracing
the algorithm's performance for a few specific inputs can be a very worthwhile
activity, it cannot prove the algorithm's correctness conclusively. But in order to
l
14
Introduction
show that an algorithm is incorrect, you need just one instance of its input for
which the algorithm fails. If the algorithm is found to be incorrect, you need to either redesign it under the same decisions regarding the data structures, the design
technique, and so on, or, in a more dramatic reversal, to reconsider one or more
of those decisions (see Figure 1.2).
The notion of correctness for approximation algorithms is less straightforward
than it is for exact algorithms. For an approximation algorithm, we usually would
like to be able to show that the error produced by the algorithm does not exceed
a predefined limit. You can find examples of such investigations in Chapter 12.
Analyzing an Algorithm
We usually want our algorithms to possess several qualities. After correctness, by
far the most important is efficiency. In fact, there are two kinds of algorithm efficiency: time efficiency and space efficiency. Time ~fficiency indicates how fast the
algorithm runs; space efficiency indicates how much extra memory the algorithm
needs. A general framework and specific techniques for analyzing an algorithm's
efficiency appear in Chapter 2.
Another desirable characteristic of an algorithm is simplicity. Unlike efficiency, which can be precisely defined and investigated with mathematical rigor,
simplicity, like beauty, is to a considerable degree in the eye of the beholder. For
example, most people would agree that Euclid's algorithm is simpler than the
middle-school procedure for computing gcd(m, n), but it is not clear whether Euclid's algorithm is simpler than the consecutive integer checking algorithm. Still,
simplicity is an important algorithm characteristic to strive for. Why? Because simpler algorithms are easier to understand and easier to program; consequently, the
resulting programs usually contain fewer bugs. There is also the undeniable aesthetic appeal of simplicity. Sometimes simpler algorithms are also more efficient
than more complicated alternatives. Unfortunately, it is not always true, in which
case a judicious compromise needs to be made.
Yet another desirable characteristic of an algorithm is generality. There are,
in fact, two issues here: generality of the problem the algorithm solves and the
range of inputs it accepts. On the first issue, note that it is sometimes easier to
design an algorithm for a problem posed in more general terms. Consider, for
example, the problem of determining whether two integers are relatively prime,
i.e., whether their only common divisor is equal to 1. It is easier to design an
algorithm for a more general problem of computing the greatest common divisor
of two integers and, to solve the former problem, check whether the gcd is 1 or
not. There are situations, however, where designing a more general algorithm is
unnecessary or difficult or even impossible. For example, it is unnecessary to sort
a list of n numbers to find its median, which is its rn/21 th smallest element. To give
another example, the standard formula for roots of a quadratic equation cannot
be generalized to handle polynomials of arbitrary degrees.
1.2 Fundamentals of Algorithmic Problem Solving
15
As to the range of inputs, your main concern should be designing an algorithm
that can handle a range of inputs that is natural for the problem at hand. For example, excluding integers equal to 1 as possible inputs for a greatest common divisor
algorithm would be quite unnatural. On the other hand, although the standard
formula for the roots of a quadratic equation holds for complex coefficients, we
would normally not implement it on this level of generality unless this capability
is explicitly required.
If you are not satisfied with the algorithm's efficiency, simplicity, or generality, you must return to the drawing board aud redesign the algorithm. In fact,
even if your evaluation is positive, it is still worth searching for other algorithmic
solutions. Recall the three different algorithms in the previous section for computing the greatest common divisor; generally, you should not expect to get the
best algorithm on the first try. At the very least, you should try to fine-tunc the
algorithm you already have. For example, we made several improvements in our
implementation of the sieve of Eratosthenes compared with its initial outline in
Section 1.1. (Can you identify them?) You will do well if you keep in mind the
following observation of Antoine de Saint-Exupery, the French writer, pilot, and
aircraft designer: "A designer knows he has arrived at perfection not when there
is no longer anything to add, but when there is no longer anything to take away." 1
Coding an Algorithm
Most algorithms are destined to be ultimately implemented as computer programs. Programming an algorithm presents both a peril and an opportunity. The
peril lies in the possibility of making the transition from an algorithm to a program either incorrectly or very inefficiently. Some influential computer scientists
strongly believe that unless the correctness of a computer program is proven with
full mathematical rigor, the program cannot be considered correct. They have developed special techniques for doing such proofs (see [Gri81]), but the power of
these techniques of formal verification is limited so far to very small programs.
As a practical matter, the validity of programs is still established by testing. Testing of computer programs is an art rather than a science, but that does not mean
that there is nothing in it to learn. Look up books devoted to testing and debugging; even more important, test and debug your program thoroughly whenever
you implement an algorithm.
Also note that throughout the book, we assume that inputs to algorithms fall
within their specified ranges and hence require no verification. When implementing algorithms as programs to be used in actual applications, you should provide
such verifications.
1.
I found this call for design simplicity in an essay collection by Jon Bentley [BenOO]; the essays deal
with a variety of issues in algorithm design and implementation, and arc justifiably titled Programming
Pearls. I wholeheartedly recommend writings of both Jon Bentley and Antoine de Saint-Exup6ry.
r
'
.
I,
!'!
16
Introduction
Of course, implementing an algorithm correctly is necessary but not sufficient:
you would not like to diminish your algorithm's power by an inefficient implementation. Modern compilers do provide a certain safety net in this regard, especially
when they are used in their code optimization mode. Still, you need to be aware
of such standard tricks as computing a loop's invariant (an expression that does
not change its value) outside the loop, collecting common subexpressions, replacing expensive operations by cheap ones, and so on. (See [Ker99] and [BenOO] for
a good discussion of code tuning and other issues related to algorithm programming.) Typically, such improvements can speed up a program only by a constant
factor, whereas a better algorithm can make a difference in running time by orders
of magnitude. But once an algorithm is selected, a 10-50% speedup may be worth
an effort.
A working program provides an additional opportunity in allowing an empirical analysis of the underlying algorithm. The analysis is based on timing the
program on several inputs and then analyzing the results obtained. We discuss the
advantages and disadvantages of this approach to analyzing algorithms in Section 2.6.
In conclusion, let us emphasize again the main lesson of the process depicted
in Figure 1.2:
As a rule, a good algorithm is a result of repeated effort and rework.
Even if you have been fortunate enough to get an algorithmic idea that seems
perfect, you should still try to see whether it can he improved.
Actually, this is good news since it makes the ultimate result so much more
enjoyable. (Yes, I did think of naming this hook The Joy of Algorithms.) On the
other hand, how does one know when to stop? In the real world, more often than
not the project's schedule or the patience of your boss will stop you. And so it
should be: perfection is expensive and in fact not always called for. Designing
an algorithm is an engineering-like activity that calls for compromises among
competing goals under the constraints of available resources, with the designer's
time being one of the resources.
I
'.
I
I
1:
In the academic world, the question leads to an interesting but usually difficult
investigation of an algorithm's optimality. Actually, this question is not about the
efficiency of an algorithm but about the complexity of the problem it solves: what
is the minimum amount of effort any algorithm will need to exert to solve the
problem in question? For some problems, the answer to this question is known.
For example, any algorithm that sorts an array by comparing values of its elements
needs about n log2 n comparisons for some arrays of size n (see Section 11.2).
But for many seemingly easy problems, such as matrix multiplication, computer
scientists do not yet have a final answer.
Another important issue of algorithmic problem solving is the question of
whether or not every problem can be solved by an algorithm. We are not talking
here about problems that do not have a solution, such as finding real roots of
~
E
i
l
1i
il
j
!
I
1.2
Fundamentals of Algorithmic Problem Solving
17
a quadratic equation with a negative discriminant. For such cases, an output
indicating that the problem does not have a solution is all we can and should
expect from an algorithm. Nor are we talking about ambiguously stated problems.
Even some unambiguous problems that must have a simple yes or no answer are
"undecidable," i.e., unsolvable by any algorithm. An important example of such
a problem appears in Section 11.3. Fortunately, a vast majority of problems in
practical computing can be solved by an algorithm.
Before leaving this section, let us be sure that you do not have the misconception-possibly caused by the somewhat mechanical nature of the diagram
of Figure 1.2-that designing an algorithm is a dull activity. There is nothing
further from the truth: inventing (or discovering?) algorithms is a very creative
and rewarding process. This book is designed to convince you that this is the case.
-----Exercises 1 . 2 - - - - - - - - - - - - - - 1. Old World puzzle A peasant finds himself on a riverbank with a wolf, a goat,
and a head of cabbage. He needs to transport all three to the other side of the
river in his boat. However, the boat has room for only the peasant himself
and one other item (either the wolf, the goat, or the cabbage). In his absence,
the wolf would eat the goat, and the goat would eat the cabbage. Solve this
problem for the peasant or prove it has no solution. (Note: The peasant is a
vegetarian but does not like cabbage and hence can eat neither the goat nor
the cabbage to help him solve the problem. And it goes without saying that
the wolf is a protected species.)
2. New World puzzle There are four people who want to cross a bridge; they
all begin on the same side. You have 17 minutes to get them all across to
the other side. It is night, and they have one flashlight. A maximum of two
people can cross the bridge at one time. Any party that crosses, either one
or two people, must have the flashlight with them. The flashlight must be
walked back and forth; it cannot be thrown, for example. Person 1 takes 1
minute to cross the bridge, person 2 takes 2 minutes, person 3 takes 5 minutes,
and person 4 takes 10 minutes. A pair must walk together at the rate of the
slower person's pace. For example, if person 1 and person 4 walk across first,
10 minutes have elapsed when they get to the other side of the bridge. If person
4 returns the flashlight, a total of 20 minutes have passed and you have failed
the mission. (Note: According to a rumor on the Internet, interviewers at a
well-known software company located near Seattle have given this problem
to interviewees.)
3. Which of the following formulas can be considered an algorithm for computing the area of a triangle whose side lengths are given positive numbers a, b,
and c?
18
Introduction
a. S = ,j p(p- a)(p- b)(p- c), where p =(a+ b + c)/2
b. S = 1hc sin A, where A is the angle between sides band c
c. S = iaha, where ha is the height to base a
4. Write a pseudocode for an algorithm for finding real roots of equation ax 2 +
bx + c = 0 for arbitrary real coefficients a, b, and c. (You may assume the
availability of the square root function sqrt(x).)
5. Describe the standard algorithm for finding the binary representation of a
positive decimal integer
a. in English.
b. in a pseudocode.
6. Describe the algorithm used by your favorite ATM machine in dispensing
cash. (You may give your description in either English or a pseudocode,
whichever you find more convenient.)
7. a. Can the problem of computing the number n be solved exactly?
b. How many instances does this problem have?
c. Look up an algorithm for this problem on the World Wide Web.
8. Give an example of a problem other than computing the greatest common
divisor for which you know more than one algorithm. Which of them is
simpler? Which is more efficient?
9. Consider the following algorithm for finding the distance between the two
closest elements in an array of numbers.
ii
i
ALGORITHM MinDistance(A[O .. n - 1])
//Input: Array A[O .. n- 1] of numbers
//Output: Minimum distance between two of its elements
dmin +-- oo
fori <--0ton-1do
for} <--0ton-1do
if i I J and IA[i]- A[J]I < dmin
dmin <-- IA[i]- A[J]I
return dmin
Make as many improvements as you can in this algorithmic solution to the
problem. (If you need to, you may change the algorithm altogether; if not,
improve the implementation given.)
10. One of the most influential books on problem solving, titled How to Solve
It [Pol57], was written by the Hungarian-American mathematician George
Polya (1887-1985). Polya summarized his ideas in a four-point summary. Find
1.3
Important Problem Types
19
this summary on the Web or. better yet. in his book. and compare it with the
plan outlined in Section 1.2. What do they have in common? How are they
different?
1.3 Important Problem Types
In the limitless sea of problems one encounters in computing, there are a few areas
that have attracted particular attention from researchers. By and large, interest
has been driven either by the problem's practical importance or by some specific
characteristics making the problem an interesting research subject; fortunately,
these two motivating forces reinforce each other in most cases.
In this section, we take up the most important problem types:
ill
"'
"
"
"
"
"'
Sorting
Searching
String processing
Graph problems
Combinatorial problems
Geometric problems
Numerical problems
These problems are used in subsequent chapters of the book to illustrate different
algorithm design techniques and methods of algorithm analysis.
Sorting
The sorting problem asks us to rearrange the items of a given list in ascending
order. Of course, for this problem to be meaningful, the nature of the list items
must allow such an ordering. (Mathematicians would say that there must exist
a relation of total ordering.) As a practical matter, we usually need to sort lists
of numbers, characters from an alphabet, character strings, and, most important,
records similar to those maintained by schools about their students, libraries about
their holdings, and companies about their employees. In the case of records, we
need to choose a piece of information to guide sorting. For example, we can choose
to sort student records in alphabetical order of names or by student number or by
student grade point average. Such a specially chosen piece of information is called
a key. Computer scientists often talk about sorting a list of keys even when the list's
items are not records but, say, just integers.
Why would we want a sorted list? Well, sorting makes many questions about
the list easier to answer. The most important of them is searching: it is why
dictionaries, telephone books, class lists, and so on are sorted. You will see other
r
'
.
'
20
Introduction
'
examples of the usefulness of list presorting in Section 6.1. In a similar vein, sorting
is used as an auxiliary step in several important algorithms in other areas, e.g.,
geometric algorithms.
By now, computer scientists have discovered dozens of different sorting algorithms. In fact, inventing a new sorting algorithm has been likened to designing
the proverbial mousetrap. And I am happy to report that the hunt for a better
sorting mousetrap continues. This perseverance is admirable in view of the following facts. On the one hand, there are a few good sorting algorithms that sort
an arbitrary array of size n using about n log2 n comparisons. On the other hand,
no algorithm that sorts by key comparisons (as opposed to, say, comparing small
pieces of keys) can do substantially better than that.
There is a reason for this embarrassment of algorithmic riches in the land
of sorting. Although some algorithms are indeed better than others, there is no
algorithm that would be the best solution in all situations. Some of the algorithms
are simple but relatively slow while others are faster but more complex; some
work better on randomly ordered inputs while others do better on almost sorted
lists; some are suitable only for lists residing in the fast memory while others can
be adapted for sorting large files stored on a disk; and so on.
Two properties of sorting algorithms deserve special mention. A sorting algorithm is called stable if it preserves the relative order of any two equal elements in
its input. In other words, if an input list contains two equal elements in positions
i and j where i < j, then in the sorted list they have to be in positions i' and j',
respectively, such that i' < j'. This property can be desirable if, for example, we
have a list of students sorted alphabetically and we want to sort it according to
student GPA: a stable algorithm will yield a list in which students with the same
GPA will still be sorted alphabetically. Generally speaking, algorithms that can
exchange keys located far apart are not stable but they usually work faster; you
will see how this general comment applies to important sorting algorithms later
in the book.
The second notable feature of a sorting algorithm is the amount of extra
memory the algorithm requires. An algorithm is said to be in place if it does
not require extra memory, except, possibly, for a few memory units. There are
important sorting algorithms that are in place and those that are not.
Searching
The searching problem deals with finding a given value, called a search key, in a
given set (or a multiset, which permits several elements to have the same value).
There are plenty of searching algorithms to choose from. They range from the
straightforward sequential search to a spectacularly efficient but limited binary
search and algorithms based on representing the underlying set in a different form
more conducive to searching. The latter algorithms are of particular importance
for real-life applications because they are indispensable for storing and retrieving
information from large databases.
1.3
Important Problem Types
21
For searching, too, there is no single algorithm that fits all situations best. Some
algorithms work faster than others but require more memory; some are very fast
but applicable only to sorted arrays; and so on. Unlike with sorting algorithms,
there is no stability problem, but different issues arise. Specifically, in applications where the underlying data may change frequently relative to the uumber
of searches, searching has to be considered in conjunction with two other operations: addition to and deletion from the data set of an item. In such situations,
data structures and algorithms should be chosen to strike a balance among the
requirements of each operation. Also, organizing very large data sets for efficient
searching poses special challenges with important implications for real-life applications.
String Processing
In recent years, the rapid proliferation of applications dealing with nonnumerical
data has intensified the interest of researchers and computing practitioners in
string-handling algorithms. A string is a sequence of characters from an alphabet.
Strings of particular interest are text strings, which comprise letters, numbers, and
special characters; bit strings, which comprise zeros and ones; and gene sequences,
which can be modeled by strings of characters from the four-character alphabet {A,
C, G, T}. It should be pointed out, however, that string-processing algorithms have
been important for computer science for a long time in conjunction with computer
languages and compiling issues.
One particular problem-that of searching for a given word in a text-has
attracted special attention from researchers. They call it string matching. Several
algorithms that exploit the special nature of this type of searching have been
invented. We introduce one very simple algorithm in Chapter 3, and discuss two
algorithms based on a remarkable idea by R. Boyer and J. Moore in Chapter 7.
Graph Problems
One of the oldest and most interesting areas in algorithmics is graph algorithms.
Informally, a graph can be thought of as a collection of points called vertices, some
of which are connected by line segments called edges. (A more formal definition
is given in the next section.) Graphs are an interesting subject to study for both
theoretical and practical reasons. Graphs can be used for modeling a wide variety
of real-life applications, including transportation and communication networks,
project scheduling, and games. One interesting recent application is an estimation
of the Web's diameter, which is the maximum number of links one needs to follow
to reach one Web page from another by the most direct route between them 2
2.
This number, according to an estimate by a group of researchers at the University of Notre Dame
[Alb99], is just 19.
22
Introduction
Basic graph algorithms include graph traversal algorithms (How can one
visit all the points in a network?), shortest-path algorithms (What is the best
route between two cities?), and topological sorting for graphs with directed edges
(Is a set of courses with their prerequisites consistent or self-contradictory?).
Fortunately, these algorithms can be considered illustrations of general design
techniques; accordingly, you will find them in corresponding chapters of the book.
Some graph problems are computationally very hard; the most well-known
examples are the traveling salesman problem and the graph-coloring problem.
The traveling salesman problem (TSP) is the problem of finding the shortest
tour through n cities that visits every city exactly once. In addition to obvious
applications involving route planning, it arises in such modern applications as
circuit board and VLSI chip fabrication, X-ray chrystallography, and genetic
engineering. The gmph-coloring problem asks us to assign the smallest number
of colors to vertices of a graph so that no two adjacent vertices are the same
color. This problem arises in several applications, such as event scheduling: if the
events are represented by vertices that are connected by an edge if and only if
the corresponding events cannot be scheduled in the same time, a solution to the
graph-coloring problem yields an optimal schedule.
Combinatorial Problems
From a more abstract perspective, the traveling salesman problem and the graphcoloring problem are examples of combinatorial problems. Tbese are problems
that ask (explicitly or implicitly) to find a combinatorial object-such as a permutation, a combination, or a subset-that satisfies certain constraints and has some
desired property (e.g., maximizes a value or minimizes a cost).
Generally speaking, combinatorial problems are the most difficult problems
in computing, from both the theoretical and practical standpoints. Their difficulty
sterns from the following facts. First, the number of combinatorial objects typically
grows extremely fast with a problem's size, reaching unimaginable magnitudes
even for moderate-sized instances. Second, there are no known algorithms for
solving most such problems exactly in an acceptable amount of time. Moreover,
most computer scientists believe that such algorithms do not exist. This conjecture
has been neither proved nor disapproved, and it remains the most important
unresolved issue in theoretical computer science. We discuss this topic in more
detail in Section 11.3.
Some combinatorial problems can be solved by efficient algorithms, but they
should be considered fortunate exceptions to the rule. The shortest -path problem
mentioned earlier is among such exceptions.
Geometric Problems
Geometric algorithms deal with geometric objects such as points, lines, and polygons. Ancient Greeks were very much interested in developing procedures (they
did not call them algorithms, of course) for solving a variety of geometric problems,
1.3
Important Problem Types
23
including problems of constructing simple geometric shapes-triangles, circles,
and so on-with an unmarked ruler and a compass. Then, for about 2000 years,
intense interest in geometric algorithms disappeared, to be resurrected in the age
o£ computers-no more rulers and compasses, just bits, bytes, and good old human ingenuity. 0£ course, today people are interested in geometric algorithms
with quite different applications in mind, such as computer graphics, robotics, and
tomography.
We will discuss algorithms for only two classic problems of computational
geometry: the closest-pair problem and the convex-hull problem. The closest-pair
problem is self-explanatory: given n points in the plane, find the closest pair among
them. The convex-hull problem asks to find the smallest convex polygon that
would include all the points of a given set. If you are interested in other geometric
algorithms, you will find a wealth of material in specialized monographs (e.g.,
[0Ro98]), or corresponding chapters of textbooks organized around problem
types (e.g., [Sed88]).
Numerical Problems
Numerical problems, another large special area of applications, are problems
that involve mathematical objects of continuous nature: solving equations and
systems of equations, computing definite integrals, evaluating functions, and so on.
The majority of such mathematical problems can be solved only approximately.
Another principal difficulty stems from the fact that such problems typically
require manipulating real numbers, which can be represented in a computer only
approximately. Moreover, a large number of arithmetic operations performed on
approximately represented numbers can lead to an accumulation of the round-off
error to a point where it can drastically distort an output produced by a seemingly
sound algorithm.
Many sophisticated algoritluns have been developed over the years in this
area, and they continue to play a critical role in many scientific and engineering
applications. But in the last 25 years or so, the computing industry has shifted
its focus to business applications. These new applications require primarily algorithms for information storage, retrieval, transportation through networks, and
presentation to users. As a result of this revolutionary change) numerical analysis
has lost its formerly dominating position in both industry and computer science
programs. Still, it is important for any computer-literate person to have at least a
rudimentary idea about numerical algorithms. We discuss several classical numer-
ical algorithms in Sections 6.2, 11.4, and 12.4.
-----Exercises 1 o3 - - - - - - - - - - - - - - - 1. Consider the algorithm for the sorting problem that sorts an array by counting,
for each of its elements, the number of smaller elements and then uses this
information to put the element in its appropriate position in the sorted array:
24
Introduction
ALGORITHM
ComparisonCountingSort(A[O .. n -1])
//Sorts an array by comparison counting
//Input: Array A[O .. n -1] of orderable values
//Output: Array S[O .. n- 1] of A's elements sorted in nondecreasing order
for i <-- 0 to n - 1 do
C ount[i] <-- 0
for i <-- 0 to n - 2 do
'
'
\..
for j +-- i + 1 to n - 1 do
':!\,.I./!
if A[i] < A[j]
Count[j] <-- Count[j] + 1
else Count[i] +-- Count[i] + 1
fori +-0ton-1do
S[Count[i]] <-- A[i]
returnS
a. Apply this algorithm to sorting the list 60, 35, 81, 98, 14, 47.
b. Is this algorithm stable?
c. Is it in place?
2. Name the algorithms for the searching problem that you already know. Give
a good succinct description of each algorithm in English. (If you know no such
algorithms, use this opportunity to design one.)
3. Design a simple algorithm for the string-matching problem.
4. Konigsberg bridges The Konigsberg bridge puzzle is universally accepted
as the problem that gave birth to graph theory. It was solved by tbe great
Swiss-born mathematician Leonhard Euler (1707-1783). The problem asked
whether one could, in a single stroll, cross all seven bridges of the city of
Konigsberg exactly once and return to a starting point. Following is a sketch
of the river with its two islands and seven bridges:
a. State the problem as a graph problem.
b. Does this problem have a solution? If you believe it does, draw such a stroll;
if you believe it does not, explain why and indicate the smallest number of
new bridges that would be required to make such a stroll possible.
II
li
!I
I
1.3
Important Problem Types
25
5. Icosian Game A century after Euler's discovery (see Problem 4), another
famous puzzle-this one invented by the renown Irish mathematician Sir
William Hamilton (1805-1865)-was presented to the world under the name
of the Icosian Game. The game was played on a circular wooden board on
which the following graph was carved:
Find a Hamiltonian circuit-a path that visits all the graph's vertices exactly
once before returning to the starting vertex-for this graph.
6. Consider the following problem: Design an algorithm to determine the best
route for a subway passenger to take from one designated station to another in
a well-developed subway system similar to those in such cities as Washington,
D.C., and London, UK.
a. The problem's statement is somewhat vague, which is typical of real-life
problems. In particular, what reasonable criterion can be used for defining
the "best" route?
b. How would you model this problem by a graph?
7. a. Rephrase the traveling salesman problem in combinatorial object terms.
b. Rephrase the graph-coloring problem in combinatorial object terms.
8. Consider the following map:
26
Introduction
a. Explain how we can use the graph-coloring problem to color the map so
that no two neighboring regions are colored the same.
b. Use your answer to part (a) to color the map with the smallest number of
colors.
9. Design an algorithm for the following problem: Given a set of n points in the
Cartesian plane, determine whether all of them lie on the same circumference.
10. Write a program that reads as its inputs the (x, y) coordinates of the endpoints
of two line segments P1 Q 1 and P2 Q2 and determines whether the segments
have a common point.
i'
,I
1 .4 Fundamental Data Structures
Since the vast majority of algorithms of interest operate on data, particular ways of
organizing data play a critical role in the design and analysis of algorithms. A data
structure can be defined as a particular scheme of organizing related data items.
The nature of the data items is dictated by a problem at hand; they can range
from elementary data types (e.g., integers or characters) to data structures (e.g., a
one-dimensional array of one-dimensional arrays is often used for implementing
matrices). There are a few data structures that have proved to be particularly
important for computer algorithms. Since you are undoubtedly familiar with most
if not all of them, just a quick review is provided here.
Linear Data Structures
The two most important elementary data structures are the array and the linked
list. A (one-dimensional) array is a sequence of n items of the same data type that
are stored contiguously in computer memory and made accessible by specifying a
value of the array's index (Figure 1.3).
In the majority of cases, the index is an integer either between 0 and n - 1
(as shown in Figure 1.3) or between 1 and n. Some computer languages allow an
array index to range between any two integer bounds low and high, and some even
permit nonnumerical indices to specify, for example, data items corresponding to
the 12 months of the year by the month names.
Item 101
Item [11
FIGURE 1.3 Array of n elements
Item [o-11
1.4
Fundamental Data Structures
27
FIGURE 1.4 Singly linked list of n elements
Each and every element of an array can be accessed in the same constant
amount of time regardless of where in the array the element in question is located.
This feature positively distinguishes arrays from linked lists (see below). It is also
assumed that every element of an array occupies the same amount of computer
storage.
Arrays are used for implementing a variety of other data structures. Prominent among them is the string, a sequence of characters from an alphabet terminated by a special character indicating the string's end. Strings composed of zeros
and ones are called binary strings or bit strings. Strings are indispensable for processing textual data, defining computer languages and compiling programs written
in them, and studying abstract computational models. Operations we usually perform on strings differ from those we typically perform on other arrays (say, arrays
of numbers). They include computing the string length, comparing two strings to
determine which one precedes the other according to the so-called lexicographic
order, i.e., in a dictionary, and concatenating two strings (forming one string from
two given strings by appending the second to the end of the first).
A linked list is a sequence of zero or more elements called nodes each
containing two kinds of information: some data and one or more links called
pointers to other nodes of the linked list. (A special pointer called "null" is used
to indicate the absence of a node's successor.) In a singly linked list, each node
except the last one contains a single pointer to the next element (Figure 1.4).
To access a particular node of a linked list, we start with the list's first node
and traverse the pointer chain until the particular node is reached. Thus, the time
needed to access an element of a singly linked list, unlike that of an array, depends
on where in the list the element is located. On the positive side, linked lists do
not require any preliminary reservation of the computer memory, and insertions
and deletions can be made quite efficiently in a linked list by reconnecting a few
appropriate pointers.
We can exploit flexibility of the linked list structure in a variety of ways. For
example, it is often convenient to start a linked list with a special node called
the header. This node often contains information about the linked list such as its
current length; it may also contain, in addition to a pointer to the first element, a
pointer to the linked list's last element.
Another extension is the structure called the doubly linked list, in which every
node, except the first and the last, contains pointers to both its successor and its
predecessor (Figure 1.5).
The array and linked list are two principal choices in representing a more
abstract data structure called a linear list or simply a list. A list is a finite sequence
28
Introduction
-lnulll/temO I Fl·l/tem11 J- ···~1·1/temn-++FIGURE 1.5 Doubly linked list of n elements
~
I
of data items, i.e., a collection of data items arranged in a certain linear order. The
basic operations performed on this data structure are searching for, inserting, and
deleting an element.
Two special types of lists, stacks and queues, are particularly important. A
stack is a list in which insertions and deletions can be done only at the end. This end
is called the top because a stack is usually visualized not horizontally but vertically
(akin to a stack of plates whose "operations" it mimics very closely). As a result,
when elements are added to (pushed onto) a stack and deleted from (popped off)
it, the structure operates in the "last-in-first-out" (LIFO) fashion, exactly as the
stack of plates does if we can remove only the top plate or add another plate to
top of the stack. Stacks have a multitude of applications; in particular, they are
indispensable for implementing recursive algorithms.
A queue, on the other hand, is a list from which elements are deleted from
one end of the structure, called the front (this operation is called dequeue),
and new elements are added to the other end, called the rear (this operation is
called enqueue). Consequently, a queue operates in the "first -in-first-out" (FIFO)
fashion (akin, say, to a queue of customers served by a single teller in a bank).
Queues also have many important applications, including several algorithms for
graph problems.
Many important applications require selection of an item of the highest priority among a dynamically changing set of candidates. A data structure that seeks to
satisfy the needs of such applications is called a priority queue. A priority queue
is a collection of data items from a totally ordered universe (most often, integer or real numbers). The principal operations on a priority queue are finding its
largest element, deleting its largest element, and adding a new element. Of course,
a priority queue must be implemented so that the last two operations yield another priority queue. Straightforward implementations of this data structure can
be based on either an array or a sorted array, but neither of these options yields
the most efficient solution possible. A better implementation of a priority queue
is based on an ingenious data structure called the heap. We discuss heaps (and an
important sorting algorithm based on them) in Section 6.4.
Graphs
As mentioned in the previous section, a graph is informally thought of as a collection of points in the plane called "vertices" or "nodes," some of them connected
by line segments called "edges" or "arcs." Formally, a graph G = {V, E) is defined
by a pair of two sets: a finite set V of items called vertices and a set E of pairs
1.4
Fundamental Data Structures
29
of these items called edges. lf these pairs of vertices are unordered, i.e., a pair of
vertices (u, v) is the same as the pair (v, u), we say that the vertices u and v are
adjacent to each other and that they are connected by the undirected edge (u, v).
We call the vertices u and v endpoints of the edge (u, v) and say that u and v are
incident to this edge; we also say that the edge (u, v) is incedent to its endpoints
u and v. A graph G is called undirected if every edge in it is undirected.
If a pair of vertices (u, v) is not the same as the pair (v, u), we say that the edge
(u, v) is directed from the vertex u, called the edge's tail, to the vertex v, called
the edge's head. We also say that the edge (u, v) leaves u and enters v. A graph
whose every edge is directed is called directed. Directed graphs are also called
digraphs.
It is normally convenient to label vertices of a graph or a digraph with letters,
integer numbers, or, if an application calls for it, character strings (Figure 1.6). The
graph in Figure 1.6a has six vertices and seven edges:
V ={a, b, c, d, e, f), E ={(a, c), (a, d), (b, c), (b, f), (c, e), (d, e), (e, f)).
The digraph in Figure 1.6b has six vertices and eight directed edges:
V ={a, b, c, d, e, f),
E ={(a, c), (b, c), (b, f), (c, e), (d, a), (d, e), (e, c), (e, f)).
Our definition of a graph does not forbid loops, or edges connecting vertices
to themselves. Unless explicitly stated otherwise, we will consider graphs without
loops. Since our definition disallows multiple edges between the same vertices of
an undirected graph, we have the following inequality for the number of edges IE I
possible in an undirected graph with IV I vertices and no loops:
0 :<:lEIs IVICIVI- 1)/2.
(We get the largest number of edges in a graph if there is an edge connecting
each of its 1V 1 vertices with all IV I - 1 other vertices. We have to divide product
IVICIVI - 1) by 2, however, because it includes every edge twice.)
A graph with every pair of its vertices connected by an edge is called complete.
A standard notation for the complete graph with lVI vertices is KIVI· A graph
with relatively few possible edges missing is called dense; a graph with few edges
(a)
FIGURE 1.6 (a) Undirected graph. (b) Digraph.
(b)
30
Introduction
a
b
c
a b c d
e
0
0
0
0
d
e
f
0
0
0
0
0
0
0
0
0
0
0
0
0
0
(a)
f
0
1
0
0
1
0
-7
c
c
-7
a
-7
d
f
b
-7
a
c
b
-7
e
-7
-7
-7
-7
-7
-7
d
-7
e
-7
e
-7
f
(b)
FIGURE 1.7 (a) Adjacency matrix and (b) adjacency lists of the graph in Figure 1.6a
relative to the number of its vertices is called sparse. Whether we are dealing with
a dense or sparse graph may influence how we choose to represent the graph and,
consequently, the running time of an algorithm being designed or used.
Graph representations Graphs for computer algorithms can be represented in
two principal ways: the adjacency matrix and adjacency lists. The adjacency matrix
of a graph with n vertices is an n-by-n boolean matrix with one row and one
column for each of the graph's vertices, in which the element in the ith row
and the jth column is equal to 1 if there is an edge from the ith vertex to the
jth vertex, and equal to 0 if there is no such edge. For example, the adjacency
matrix for the graph in Figure 1.6a is given in Figure 1.7a. Note that the adjacency
matrix of an undirected graph is always symmetric, i.e., A[i, j] = A[j, i] for every
0 :0 i, j :0 n - 1 (why?).
The adjacency lists of a graph or a digraph is a collection of linked lists,
one for each vertex, that contain all the vertices adjacent to the list's vertex
(i.e., all the vertices connected to it by an edge). Usually, such lists start with a
header identifying a vertex for which the list is compiled. For example, Figure 1.7b
represents the graph in Figure 1.6a via its adjacency lists. To put it another way,
adjacency lists indicate columns of the adjacency matrix that, for a given vertex,
contain 1's.
If a graph is sparse, the adjacency list representation may use less space
than the corresponding adjacency matrix despite the extra storage consumed by
pointers of the linked lists; the situation is exactly opposite for dense graphs. In
general, which of the two representations is more convenient depends on the
nature of the problem, on the algorithm used for solving it, and, possibly, on the
type of input graph (sparse or dense).
Weighted graphs A weighted graph (or weighted digraph) is a graph (or digraph)
with numbers assigned to its edges. These numbers are called weights or costs. An
interest in such graphs is motivated by numerous real-life applications, such as
finding the shortest path between two points in a transportation or communication
network or the traveling salesman problem mentioned earlier.
1.4
a
!~ r
1
4
2
(a)
b
c
5
1
d
00
31
Fundamental Data Structures
c
b
5
d
---7b,5---7C,1
00
7
7
00
4
2
(b)
I] m
--7
a, 5 ---7 c, 7 -;. d, 4
-> a, 1 -. b, 7 -. d, 2
-;> b, 4 -;> C, 2
(c)
FIGURE 1.8 (a) Weighted graph. (b) Its weight matrix. (c) Its adjacency lists.
Both principal representations of a graph can be easily adopted to accommodate weighted graphs. If a weighted graph is represented by its adjacency matrix,
then its element A[i, j] will simply contain the weight of the edge from the ith to
the jth vertex if there is such an edge and a special symbol, e.g., co, if there is no
such edge. Such a matrix is called the weight matrix or cost matrix. This approach
is illustrated in Figure l.Sb. (For some applications, it is more convenient to put
O's on the main diagonal of the adjacency matrix.) Adjacency lists for a weighted
graph have to include in their nodes not only the name of an adjacent vertex but
also the weight of the corresponding edge (Figure 1.8c).
Paths and cycles Among many interesting properties of graphs, two are important for a great number of applications: connectivity and acyclicity. Both are based
on the notion of a path. A path from vertex u to vertex v of a graph G can be detined as a sequence of adjacent (connected by an edge) vertices that starts with u
and ends with v. If all vertices of a path are distinct, the path is said to be simple.
The length of a path is the total number of vertices in a vertex sequence defining
the path minus one, which is the same as the number of edges in the path. For example, a, c, b, f is a simple path oflength 3 from a to fin the graph of Figure 1.6a,
whereas a, c, e, c, b, f is a path (not simple) oflength 5 from a to f.
In the case of a directed graph, we are usually interested in directed paths.
A directed path is a sequence of vertices in which every consecutive pair of the
vertices is connected by an edge directed from the vertex listed first to the vertex
listed next. For example, a, c, e, f is a directed path from a to fin the graph of
Figure 1.6b.
A graph is said to be connected if for every pair of its vertices u and v there is
a path from u to v. Informally, this property means that if we make a model of a
connected graph by connecting some balls representing the graph's vertices with
strings representing the edges, it will be a single piece. If a graph is not connected,
such a model will consist of several connected pieces that are called connected
components of the graph. Formally, a connected component is a maximal (not
32
Introduction
f
g
h
FIGURE 1.9 Graph that is not connected
expandable via an inclusion of an extra vertex) connected subgraph3 of a given
graph. For example, the graphs of Figures 1.6a and l.Sa are connected, while the
graph in Figure 1.9 is not because there is no path, for example, from a to f. The
graph in Figure 1.9 has two connected components with vertices (a, b, c, d, e) and
(f, g, h, i], respectively.
Graphs with several connected components do happen in real-life applications. A graph representing the Interstate highway system of the United States
would be an example (why?).
It is important to know for many applications whether or not a graph under
consideration has cycles. A cycle is a path of a positive length that starts and ends at
the same vertex and does not traverse the same edge more than once. For example,
j, h, i, g, .f is a cycle in the graph of Figure 1.9. A graph with no cycles is said to
be acyclic. We discuss acyclic graphs in the next subsection.
Trees
A tree (more accurately, afree tree) is a connected acyclic graph (Figure l.lOa).
A graph that has no cycles but is not necessarily connected is called a forest: each
of its connected components is a tree (Figure l.lOb ).
Trees have several important properties other graphs do not have. In particular, the number of edges in a tree is always one less than the number of its
vertices:
lEI= IVI-1.
As the graph of Figure 1.9 demonstrates, this property is necessary but not sufficient for a graph to be a tree. However, for connected graphs it is sufficient and
hence provides a convenient way of checking whether a connected graph has a
cycle.
3.
Asubgraph of a given graph G = (V, E) is a graph G' = (V', E'} such that V' ~VandE'£;; E.
1.4
Fundamental Data Structures
(a)
33
(b)
FIGURE 1.10 (a) Tree. (b) Forest.
e
f
h
(a)
(b)
FIGURE 1.11 (a) Free tree. (b) Its transformation into a rooted tree.
Rooted trees Another very important property of trees is the fact that for every
two vertices in a tree, there always exists exactly one simple path from one of these
vertices to the other. This property makes it possible to select an arbitrary vertex
in a free tree and consider it as the root of the so-called rooted tree. A rooted tree
is usually depicted by placing its root on the top (level 0 of the tree), the vertices
adjacent to the root below it (Ievell), the vertices two edges apart from the root
below that (level2), and so on. Figure 1.11 presents such a transformation from a
free tree to a rooted tree.
Rooted trees play a very important role in computer science, a much more
important one than free trees do; in fact, for the sake of brevity, they are often
referred to as simply "trees." Obvious applications of trees are for describing
hierarchies, from file directories to organizational charts of enterprises. There are
many less obvious applications, such as implementing dictionaries (see below),
efficient storage of very large data sets (Section 7.4), and data encoding (Section
9.4). As we discuss in Chapter 2, trees also are helpful in analysis of recursive
algorithms. To finish this far-from-complete list of tree applications, we should
34
Introduction
mention the so-called state-space trees that underline two important algorithm
design techniques: backtracking and branch-and-bound (Sections 12.1 and 12.2).
For any vertex v in a tree T, all the vertices on the simple path from the root to
that vertex are called ancestors of v. The vertex itself is usually considered its own
ancestor; the set of ancestors that excludes the vertex itself is referred to as proper
ancestors. If (u, v) is the last edge of the simple path from the root to vertex v (and
u ,P v), u is said to be the parent of v and vis called a child of u; vertices that have
the same parent are said to be siblings. A vertex with no children is called a leaf; a
vertex with at least one child is called parental. All the vertices for which a vertex
v is an ancestor are said to be descendants of v; the proper descendants exclude
the vertex v itself. All the descendants of a vertex v with all the edges connecting
them form the subtree ofT rooted at that vertex. Thus, for the tree of Figure 1.11b,
the root of the tree is a; vertices d, g, .f, h, and i are leaves, while vertices a, b, e,
and c are parental; the parent of his a; the children of b are c and g; the siblings
of bared and e; the vertices of the subtree rooted at bare {b. c, g, h, i).
The depth of a vertex vis the length of the simple path from the root to v. The
height of a tree is the length of the longest simple path from the root to a leaf. For
example, the depth of vertex c in the tree in Figure 1.1lb is 2, and the height of
the tree is 3. Thus, if we count tree levels top down starting with 0 for the root's
level, the depth of a vertex is simply its level in the tree, and the tree's height is the
maximum level of its vertices. (You should be alert to the fact that some authors
define the height of a tree as the number of levels in it; this makes the height of
a tree larger by 1 than the height defined as the length of the longest simple path
from the root to a leaf.)
Ordered trees An ordered tree is a rooted tree in which all the children of each
vertex are ordered. It is convenient to assume that in a tree's diagram, all the
children are ordered left to right. A binary tree can be defined as an ordered tree
in which every vertex has no more than two children and each child is designated
as either a left child or a right child of its parent. The subtree with its root at
the left (right) child of a vertex is called the left (right) subtree of that vertex. An
example of a binary tree is given in Figure 1.12a.
5
7
10
4
(a)
,
'i
[!
'
!i
!I
FIGURE 1.12 (a) Binary tree. (b) Binary search tree.
(b)
1··
.•
1.4
Fundamental Data Structures
35
In Figure 1.12b, some numbers are assigned to vertices of the binary tree in
Figure 1.12a. Note that a number assigned to each parental vertex is larger than all
the numbers in its left subtree and smaller than all the numbers in its right subtree.
Such trees are called binary search trees. Binary trees and binary search trees have
a wide variety of applications in computer science; you will encounter some of
them throughout the book. In particular, binary search trees can be generalized
to more general kinds of search trees called multiway sew·ch trees, which are
indispensable for efficient storage of very large files on disks.
As you will see later in the book, the efficiency of most important algorithms
for binary search trees and their extensions depends on the tree's height. Therefore, the following inequalities for the height h of a binary tree with n nodes are
especially important for analysis of such algorithms:
Llog2 nj :5: h :5: n- 1.
A binary tree is usually implemented for computing purposes by a collection
of nodes corresponding to vertices of the tree. Each node contains some information associated with the vertex (its name or some value assigned to it) and two
pointers to the nodes representing the left child and right child of the vertex, respectively. Figure 1.13 illustrates such an implementation for the binary search
tree in Figure 1.12b.
A computer representation of an arbitrary ordered tree can be done by simply
providing a parental vertex with the number of pointers equal to the number of
its children. This representation may prove to be inconvenient if the number of
children varies widely among the nodes. We can avoid this inconvenience by using
nodes with just two pointers, as we did for binary trees. Here, however, the left
pointer will point to the first child of the vertex, while the right pointer will point
FIGURE 1.13 Standard implementation of the binary search tree in Figure 1 12b
r
":"
l
36
Introduction
I
I
I
t
I
'
a
null
b
null
d
c
o-t--->1 null
h
e
g
f
(a)
(b)
FIGURE 1.14 {a) First child-next sibling representation of the graph in Figure 1.11b.
{b) Its binary tree representation.
to its next sibling. Accordingly, this representation is called the first child-next
sibling representation. Thus, all the siblings of a vertex are linked (via the nodes'
right pointers) in a singly linked list, with the first element of the list pointed
to by the left pointer of their parent. Figure 1.14a illustrates this representation
for the tree in Figure l.llb. It is not difficult to see that this representation
effectively transforms an ordered tree into a binary tree said to be associated with
the ordered tree. We get this representation by "rotating" the pointers about 45
degrees clockwise (see Figure 1.14b).
Sets and Dictionaries
The notion of a set plays a central role in mathematics. A set can be described
as an unordered collection (possibly empty) of distinct items called elements of
the set. A specific set is defined either by an explicit listing of its elements (e.g.,
S = {2, 3, 5, 7}) or by specifying a property that all the set's elements and only they
must satisfy (e.g., S = {n: n is a prime number and n < 10}). The most important set
operations are checking membership of a given item in a given set (whether a given
item is among the elements of the set), finding the union of two sets (which set
comprises all the elements that belong to either of the two sets or to both of them),
and finding the intersection of two sets (which set comprises all the elements that
belong to both sets).
Sets can be implemented in computer applications in two ways. The first
considers only sets that are subsets of some large set U, called the universal set.
I
.·,:.•·_.
.1.
..
1.4
Fundamental Data Structures
37
If set U has 11 elements, then any subset S of U can be represented by a bit
string of size 11, called a bit vector, in which the ith element is 1 if and only if
the ith element of U is included in set S. Thus, to continue with our example, if
U = (1, 2, 3, 4, 5, 6, 7, 8, 9], then S = (2, 3, 5, 7] will be represented by the bit
string 011010100. This way of representing sets makes it possible to implement
the standard set operations very fast but at the expense of potentially using a large
amount of storage.
The second and more common way to represent a set for computing purposes
is to use the list structure to indicate the set's elements. (Of course, this option,
too, is feasible only for finite sets; fortunately, unlike mathematics, this is the kind
of sets most computer applications need.) Note, however, the two principal points
of distinction between sets and lists. First, a set cannot contain identical elements;
a list can. This requirement for uniqueness is sometimes circumvented by the
introduction of a multiset or a bag, an unordered collection of items that are not
necessarily distinct. Second, a set is an unordered collection of items; therefore,
changing the order of its elements does not change the set. A list, defined as an
ordered collection of items, is exactly the opposite. This is an important theoretical
distinction, but fortunately it is not important for many applications. It is also
worth mentioning that if a set is represented by a list, depending on the application
at hand, it might be worth maintaining the list in a sorted order.
In computing, the operations we need to perform for a set or a multiset most
often are searching for a given item, adding a new item, and deleting an item
from the collection. A data structure that implements these three operations is
called the dictionary. Note the relationship between this data structure and the
problem of searching mentioned in Section 1.3; obviously, we are dealing here
with searching in a dynamic context. Consequently, an efficient implementation
of a dictionary has to strike a compromise between the efficiency of searching and
the efficiencies of the other two operations. There are quite a few ways a dictionary
ean be implemented. They range from an unsophisticated use of arrays (sorted or
not) to much more sophisticated techniques such as hashing and balanced search
trees, which we discuss later in the book.
A number of applications in computing require a dynamic partition of some
n-element set into a collection of disjoints subsets. After being initialized as a
collection of n one-element subsets, the collection is subjected to a sequence of
intermixed union and search operations. This problem is called the set union
problem. We discuss efficient algorithmic solutions to this problem in Section 9.2
in conjunction with one of its most important applications.
You may have noticed that in our review of basic data structures we almost
always mentioned specific operations that are typically performed for the structure in question. This intimate relationship between data and operations has been
recognized by computer scientists for a long time. It has led them in particular
to the idea of an abstmct data type (ADT): a set of abstract objects representing data items with a collection of operations that can be performed on them.
As illustrations of this notion, reread, say, our definitions of priority queue and
38
1
Introduction
dictionary. Although abstract data types could be implemented in older procedurallanguages such as Pascal (see, e.g., [Aho83]), it is much more convenient to do
so in object -oriented languages, such as C++ and Java, that support abstract data
types by means of classes.
-----Exercises 1 . 4 - - - - - - - - - - - - - - - -
I
I
il
1. Describe how one can implement each of the following operations on an array
so that the time it takes does not depend on the array's size n.
a. Delete the ith element of an array (1 sis n).
b. Delete the ith element of a sorted array (the remaining array has to stay
sorted, of course).
l1!
2. If you have to solve the searching problem for a list of n numbers, how can you
take advantage of the fact that the list is known to be sorted? Give separate
answeTs for
a. lists represented as arrays.
b. lists represented as linked lists.
3. a. Show the stack after each operation of the following sequence that starts
with the empty stack:
push(a), push(b), pop, push(c), push(d), pop
b. Show the queue after each operation of the following sequence that starts
with the empty queue:
enqueue(a), enqueue(b), dequeue, enqueue(c), enqueue(d), dequeue
4. a. Let A be the adjacency matrix of an undirected graph. Explain what property of the matrix indicates that
i. the graph is complete.
ii. the graph has a loop, i.e., an edge connecting a vertex to itself.
iii. the graph has an isolated vertex, i.e., a vertex with no edges incident
to it.
b. Answer the same questions for the adjacency list representation.
5. Give a detailed description of an algorithm for transforming a free tree into
a tree rooted at a given vertex of the free tree.
6. Prove the inequalities that bracket the height of a binary tree with n vertices:
Llog2 n J s h s n - 1.
1:
7. Indicate how the ADT priority queue can be implemented as
I
I
.
..
--,
Summary
39
a. an (unsorted) array.
b. a sorted array.
c. a binary search tree.
8. How would you implement a dictionary of a reasonably small size n if you
knew that all its elements are distinct (e.g., names of 50 states of the United
States)? Specify an implementation of each dictionary operation.
9. For each of the following applications, indicate the most appropriate data
structure:
a. answering telephone calls in the order of their known priorities
b. sending backlog orders to customers in the order they have been received
c. implementing a calculator for computing simple arithmetical expressions
10. Anagram checking Design an algorithm for checking whether two given
words are anagrams, i.e., whether one word can be obtained by permuting
the letters of the other. (For example, the words tea and eat are anagrams.)
SUMMARY
il
An algorithm is a sequence of nonarnbiguous instructions for solving a
problem in a fmite amount of time. An input to an algorithm specifies an
instance o[ the problem the algorithm solves.
"
Algorithms can be specified in a natural language or a pseudocode; they can
also be implemented as computer programs.
"
Among several ways to classify algorithms, the two principal alternatives are:
to group algorithms according to types of problems they solve;
- to group algorithms according to underlying design techniques they are
based upon.
II
The important problem types are sorting, searching, string processing, graph
problems, combinatmial problems, geometric problems, and numerical
problems.
"
Algorithm design techniques (or "strategies" or "paradigms") are general
approaches to solving problems algorithmically, applicable to a variety of
problems from different areas of computing.
"
Although designing an algorithm is undoubtedly a creative activity, one can
identify a sequence of interrelated actions involved in such a process. They
are summarized in Figure 1.2.
11
A good algorithm is usually a result of repeated efforts and rework.
40
Introduction
"
The same problem can often be solved by several algorithms. For example,
three algorithms were given for computing the greatest common divisor of
two integers: Euclid's algorithm, the consecutive integer checking algorithm,
and the middle-school algorithm (enhanced by the sieve of Eratosthenes for
generating a list of primes).
"
Algorithms operate on data. This makes the issue of data structuring critical
for efficient algorithmic problem solving. The most important elementary data
structures are the array and the linked list. They are used for representing
more abstract data structures such as the list, the stack, the queue, the graph
(via its adjacency matrix or adjacency lists), the binary tree, and the set.
"
An abstract collection of objects with several operations that can be performed on them is called an abstract data type (ADT). The list, the stack, the
queue, the priority queue, and the dictionary are important examples of abstract data types. Modem object-oriented languages support implementation
of ADTs by means of classes.
Fundamentals of
the Analysis of
Algorithm Efficiency
I often say that when you can measure what you are speaking about and
express it in numbers you know something about it; but when you cannot express it in numbers your knowledge is a meagre and unsatisfactory
kind: it may be the beginning of knowledge but you have scarcely, in your
thoughts, advanced to the stage of science, whatever the matter may be.
-Lord Kelvin {1824-19071
Not everything that can be counted counts, and not everything that counts
can be counted.
-Albert Einstein {1879-1955)
T
his chapter is devoted to analysis of algorithms. The American Heritage Dictionary defines "analysis" as "the separation of an intellectual or substantial
whole into its constituent parts for individual study." Accordingly, each of the principal dimensions of an algorithm pointed out in Section 1.2 is both a legitimate and
desirable subject of study. But the term "analysis of algorithms" is usually used in
a narrower technical sense to mean an investigation of an algorithm's efficiency
with respect to two resources: running time and memory space. This emphasis on
efficiency is easy to explain. First, unlike such dimensions as simplicity and generality, efficiency can be studied in precise quantitative terms. Second, one can
argue-although this is hardly always the case, given the speed and memory of
today's computers-that the efficiency considerations are of primary importance
from the practical point of view. In this chapter, we too limit the discussion to an
algorithm's efficiency.
41
rr·
!
42
Fundamentals of the Analysis of Algorithm Efficiency
We start with a general framework for analyzing algorithm efficiency in Section 2.1. This section is arguably the most important in the chapter; the fundamental nature of the topic makes it also one of the most important sections in the
entire book.
In Section 2.2, we introduce three notations: 0 ("big oh"), Q ("big omega"),
and 19 ("big theta"). Borrowed from mathematics, these notations have become
the language for discussing an algorithm's efficiency.
In Section 2.3, we show how the general framework outlined in Section 2.1 can
be systematically applied to analyzing the efficiency of nourecursive algorithms.
The main tool of such an analysis is setting up a sum representing the algorithm's
running time and then simplifying the sum by using standard sum manipulation
techniques.
In Section 2.4, we show how the general framework outlined in Section 2.1
can be systematically applied to analyzing the efficiency of recursive algorithms.
Here, the main tool is not a sum but a special kind of equation called a recurrence relation. We explain how such recurrence relations can be set up and then
introduce a method for solving them.
Although we illustrate the analysis framework and the methods of its applications by a variety of examples in the first four sections of this chapter, Section
2.5 is devoted to yet another example-that of the Fibonacci numbers. Introduced 800 years ago, this remarkable sequence appears in a variety of applications
both within and outside computer science. A discussion of the Fibonacci sequence
serves as a natural vehicle for introducing an important class of recurrence relations not solvable by the method of Section 2.4. We also discuss several algorithms
for computing the Fibonacci numbers, mostly for the sake of a few general observations about the efficiency of algorithms and methods of analyzing them.
The methods of Sections 2.3 and 2.4 provide a powerful technique for analyzing the efficiency of many algorithms with mathematical clarity and precision, but
these methods are far from being foolproof. The last two sections of the chapter
deal with two approaches-empirical analysis and algorithm visualization-that
complement the pure mathematical techniques of Sections 2.3 and 2.4. Much
newer and, hence, less developed than their mathematical counterparts, these approaches promise to play au important role among the tools available for analysis
of algorithm efficiency.
2.1
l
II
I
I
~
l'
~
I'
Analysis Framework
In this section, we outline a general framework for analyzing the efficiency of algorithms. To begin with, there are two kinds of efficiency: time efficiency and space
efficiency. Time efficiency indicates how fast an algorithm in question runs; space
efficiency deals with the extra space the algorithm requires. In the early days of
electronic computing, both resources-time and space-were at a premium. Half
II
i]!l
,,Ui
2.1
Analysis Framework
43
a century of relentless technological innovations have improved the computer's
speed and memory size by many orders of magnitude. Now the amount of extra
space required by an algorithm is typically not of as much concern, with the caveat
that there is still, of course, a difference between the fast main memory, the slower
secondary memory, and the cache. The time issue has not diminished quite to the
same extent, however. In addition, the research experience has shown that for
most problems, we can achieve much more spectacular progress in speed than in
space. Therefore, following a well-established tradition of algorithm textbooks, we
primarily concentrate on time efficiency, but the analytical framework introduced
here is applicable to analyzing space efficiency as well.
Measuring an Input's Size
Let us start with the obvious observation that almost all algorithms run longer
on larger inputs. For example, it takes longer to sort larger arrays, multiply larger
matrices, and so on. Therefore, it is logical to investigate an algorithm's efficiency
as a function of some parameter n indicating the algorithm's input size. 1 In most
cases, selecting such a parameter is quite straightforward. For example, it will be
the size of the list for problems of sorting, searching, finding the list's smallest
element, and most other problems dealing with lists. For the problem of evaluating
a polynomial p(x) = a,x" + · · · + a 0 of degree n, it will be the polynomial's degree
or the number of its coefficients, which is larger by one than its degree. You will
see from the discussion that such a minor difference is inconsequential for the
efficiency analysis.
There are situations, of course,. where the choice of a parameter indicating
an input size does matter. One such example is computing the product of two nby-n matrices. There are two natural measures of size for this problem. The first
and more frequently used is the matrix order n. But the other natural contender
is the total number of elements N in the matrices being multiplied. (The latter
is also more general since it is applicable to matrices that are not necessarily
square.) Since there is a simple formula relating these two measures, we can easily
switch from one to the other, but the answer about an algorithm's efficiency will
be qualitatively different depending on which of the two measures we use (see
Problem 2 in Exercises 2.1).
The choice of an appropriate size metric can be influenced by operations of
the algorithm in question. For example, how should we measure an input's size
for a spell-checking algorithm? If the algorithm examines individual characters of
its input, then we should measure the size by the number of characters; if it works
by processing words, we should count their number in the input.
1.
Some algorithms require more than one parameter to indicate the size of their inputs (e.g., the number
of vertices and the number of edges for algorithms on graphs represented by adjacency lists).
Fundamentals of the Analysis of Algorithm Efficiency
44
We should make a special note about measuring the size of inputs for algorithms involving properties of numbers (e.g., checking whether a given integer n
is prime). For such algorithms, computer scientists prefer measuring size by the
number b of bits in then's binary representation:
b= Llog2 nJ +1.
(2.1)
This metric usually gives a better idea about the efficiency of algorithms in question.
Units for Measuring Running lime
The next issue concerns units for measuring an algorithm's running time. Of
course, we can simply use some standard unit of time measurement-a second,
a millisecond, and so on-to measure the running time of a program implementing the algorithm. There are obvious drawbacks to such an approach, however:
dependence on the speed of a particular computer, dependence on the quality of
a program implementing the algorithm and of the compiler used in generating the
machine code, and the difficulty of clocking the actual running time of the program. Since we are after a measure of an algorithm's efficiency, we would like to
have a metric that does not depend on these extraneous factors.
One possible approach is to count the number of times each of the algorithm's
operations is executed. This approach is both excessively difficult and, as we
shall see, usually unnecessary. The thing to do is to identify the most important
operation of the algorithm, called the basic operation, the operation contributing
the most to the total running time, and compute the number of times the basic
operation is executed.
As a rule, it is not difficult to identify the basic operation of an algorithm:
it is usually the most time-consuming operation in the algorithm's innermost
loop. For example, most sorting algorithms work by comparing elements (keys)
of a list being sorted with each other; for such algorithms, the basic operation
is a key comparison. As another example, algorithms for matrix multiplication
and polynomial evaluation require two arithmetic operations: multiplication and
addition. On most computers, multiplication of two numbers takes longer than
addition, making the former an unquestionable choice for the basic operation 2
Thus, the established framework for the analysis of an algorithm's time efficiency suggests measuring it by counting the number of times the algorithm's
basic operation is executed on inputs of size n. We will find out how to compute
such a count for nonrecursive and recursive algorithms in Sections 2.3 and 2.4,
respectively.
2.
On some computers based on the so-called RISC architecture, it is not necessarily the case (see, for
example, the timing data provided by Kernighan and Pike [Ker99], pp. 185-186).
I
2.1
Analysis Framework
45
Here is an important application. Let c0 P be the execution time of an algorithm's basic operation on a particular computer, and let C(n) be the number of
times this operation needs to be executed for this algorithm. Then we can estimate
the running time T (n) of a program implementing this algorithm on that computer
by the formula
T(n) "'c0 pC(n).
Of course, this formula should be used with caution. The count C (n) does not
contain any information about operations that are not basic, and, in fact, the
count itself is often computed only approximately. Further, the constant c0 , is
also an approximation whose reliability is not always easy to assess. Still, unless
n is extremely large or very small, the formula can give a reasonable estimate of
the algorithm's running time. It also makes it possible to answer such questions as
"How much faster would this algorithm run on a machine that is ten times faster
than the one we have?" The answer is, obviously, ten times. Or, assuming that
C(n) = ~n(n- 1), how much longer will the algorithm run if we double its input
size? The answer is about four times longer. Indeed, for all but very small values
ofn,
'
1
2
12
2
1
2
12
2
C(n) = -n(n -1) = -n - -n"' -n
and therefore
T(2n)
--"'
T(n)
C0
,C(2n)
C0 pC(n)
~(2n) 2
" ' - - =4
~n2
·
Note that we were able to answer the last question without actually knowing
the value of C0 P: it was neatly cancelled out in the ratio. Also note that ~. the
multiplicative constant in the formula for the count C(n), was also cancelled out.
It is for these reasons that the efficiency analysis framework ignores multiplicative
constants and concentrates on the count's order of growth to within a constant
multiple for large-size inputs.
Orders of Growth
Why this emphasis on the count's order of growth for large input sizes? A difference in running times on small inputs is not what really distinguishes efficient
algorithms from inefficient ones. When we have to compute, for example, the
greatest common divisor of two small numbers, it is not innnediately clear how
much more efficient Euclid's algorithm is compared to the other two algorithms
discussed in Section 1.1 or even why we should care which of them is faster and
by how much. It is only when we have to find the greatest common divisor of two
large numbers that the difference in algorithm efficiencies becomes both clear and
important. For large values of n, it is the function's order of growth that counts: just
1fT
~~~ .·
46
Fundamentals of the Analysis of Algorithm Efficiency
TABLE 2.1 Va!ues (some approximate) of several functions important for
analysis of algorithms
n
log2 n
11
11log2 n
112
113
2"
n!
10
102
10 3
3.3
6.6
101
10'
103
104
105
3.3-10 1
6.6-102
1.0-104
1.3-105
1.7-106
102
103
106
109
1012
1015
103
1.3·1030
3.6-106
10 6
2.0·107
w•
105
106
10
13
17
20
to•
10 6
108
1010
1012
9.3·10157
1018
look at Table 2.1, which contains values of a few functions particularly important
for analysis of algorithms.
The magnitude of the numbers in Table 2.1 has a profound significance for
the analysis of algorithms. The function growing the slowest among these is the
logarithmic function. It grows so slowly, in fact, that we should expect a program
implementing an algorithm with a logarithmic basic-operation count to run practicallyinstantaneously on inputs of all realistic sizes. Also note that although specific
values of such a count depend, of course, on the logarithm's base, the formula
log, 11 = log, h log6 n
makes it possible to switch from one base to another, leaving the count logarithmic
but with a new multiplicative constant. This is why we omit a logarithm's base and
write simply log n in situations where we are interested just in a function's order
of growth to within a multiplicative constant.
On the other end of the spectrum are the exponential function 2" and the
factorial function n! Both these functions grow so fast that their values become
astronomically large even for rather small values of n. (This is the reason why
we did not include their values for n > 102 in Table 2.1.) For example, it would
take about 4. 1010 years for a computer making one trillion (1012) operations per
second to execute 2 100 operations. Though this is incomparably faster than it would
have taken to execute 100! operations, it is still longer than 4.5 billion (4.5 · 109 )
years-the estimated age of the planet Earth. There is a tremendous difference
between the orders of growth of the functions 2" and n !, yet both are often referred
to as "exponential-growth functions" (or simply "exponential") despite the fact
that, strictly speaking, only the former should be referred to as such. The bottom
Hne, which is important to remember, is this:
Algorithms that require an exponential number of operations are practical
for solving only problems of very small sizes.
2.1
Analysis Framework
47
Another way to appreciate the qualitative difference among the orders of
growth of the functions in Table 2.1 is to consider how they react to, say, a
twofold increase in the value of their argument n. The function log2 n increases in
value by just 1 (because log2 2n = log2 2 + log2 n = 1 + log2 n ); the linear function
increases twofold; the "n-log-n" function n log2 n increases slightly more than
twofold; the quadratic function n 2 and cubic function n 3 increase fourfold and
eightfold, respectively (because (2n) 2 = 4n 2 and (2n) 3 = 8n 3); the value of 2" gets
squared (because 22" = (2") 2); and n! increases much more than that (yes, even
mathematics refuses to cooperate to give a neat answer for n!).
Worst-Case, Best-Case, and Average-Case Efficiencies
In the beginning of this section, we established that it is reasonable to measure
an algorithm's efficiency as a function of a parameter indicating the size of the
algorithm's input. But there are many algorithms for which running time depends
not only on an input size but also on the specifics of a particular input. Consider,
as an example, sequential search. This is a straightforward algorithm that searches
for a given item (some search key K) in a list of n elements by checking successive
elements of the list until either a match with the search key is found or the list
is exhausted. Here is the algorithm's pseudocode, in which, for simplicity, a list is
implemented as an array. (It also assumes that the second condition A[i] # K will
not be checked if the first one, which checks that the array's index does not exceed
its upper bound, fails.)
ALGORITHM
SequentialSearch(A[O .. n - 1], K)
//Searches for a given value in a given array by sequential search
//Input: An array A[O .. n -1] and a search key K
//Output: The index of the frrst element of A that matches K
II
or -1 if there are no matching elements
i+-0
while i <nand A[i] # K do
i+-i+1
ifi < n return i
else return -1
Clearly, the running time of this algorithm can be quite different for the
same list size n. In the worst case, when there are no matching elements or
the frrst matching element happens to be the last one on the list, the algorithm
makes the largest number of key comparisons among all possible inputs of size n:
cworst(n) = n.
The worst-case efficiency of an algorithm is its efficiency for the worst-case
input of size n, which is an input (or inputs) of size n for which the algorithm
runs the longest among all possible inputs of that size. The way to determine
48
I
I'''
I'
il:
!j~--
Fundamentals of the Analysis of Algorithm Efficiency
the worst-case efficiency of an algorithm is, in principle, quite straightforward: we
analyze the algorithm to see what kind of inputs yield the largest value of the basic
operation's count C(n) among all possible inputs of size nand then compute this
worst-case value Cworst(n). (For sequential search, the answer was obvious. The
methods for handling less trivial situations are explained in subsequent sections of
this chapter.) Clearly, the worst-case analysis provides very important information
about an algorithm's efficiency by bounding its running time from above. In other
words, it guarantees that for any instance of size n, the running time will not exceed
Cworst(n), its running time on the worst-case inputs.
The best-case efficiency of an algorithm is its efficiency for the best-case input
of size n, which is an input (or inputs) of size n for which the algorithm runs the
fastest among all possible inputs of that size. Accordingly, we can analyze the bestcase efficiency as follows. First, we determine the kind of inputs for which the count
C(n) will be the smallest among all possible inputs of size n. (Note that the best
case does not mean the smallest input; it means the input of size n for which the
algorithm runs the fastest.) Then we ascertain the value of C (n) on these most
convenient inputs. For example, for sequential search, best-case inputs are lists of
size n with their first elements equal to a search key; accordingly, Cbw(n) = 1.
The analysis of the best-case efficiency is not nearly as important as that of the
worst-case efficiency. But it is not completely useless, either. Though we should
not expect to get best-case inputs, we might be able to take advantage of the fact
that for some algorithms a good best -case performance extends to some useful
types of inputs close to being the best-case ones. For example, there is a sorting
algorithm (insertion sort) for which the best-case inputs are already sorted arrays
on which the algorithm works very fast. Moreover, this good best-case efficiency
deteriorates only slightly for almost sorted arrays. Thus, such an algorithm might
well be the method of choice for applications dealing with almost sorted arrays.
And, of course, if the best-case efficiency of an algorithm is unsatisfactory, we can
immediately discard it without further analysis.
It should be clear from our discussion, however, that neither the worst-case
analysis nor its best -case counterpart yields the necessary information about
an algorithm's behavior on a "typical" or "random" input. This is the information that the average-case efficiency seeks to provide. To analyze the algorithm's
average-case efficiency, we must make some assumptions about possible inputs of
size n.
Let us consider again sequential search. The standard assumptions are that
(a) the probability of a successful search is equal top (0 ::0 p ::0 1) and (b) the
probability of the first match occurring in the ith position of the list is the same
for every i. Under these assumptions-the validity of which is usually difficult to
verify, their reasonableness notwithstanding-we can find the average number
of key comparisons Ca, 8 (n) as follows. In the case of a successful search, the
probability of the first match occurring in the i th position of the list is p j n for
every i, and the number of comparisons made by the algorithm in such a situation
2.1
Analysis Framework
49
is obviously i. In the case of an unsuccessful search, the number of comparisons
is n with the probability of such a search being (1- p). Therefore,
Cav (n)=[l· [!_ +2· [!_+·· ·+i · [!_ +·· ·+n · f]+n· (1- p)
g
n
n
n
n
= f[1 + 2 + · ·. + i + · · · + n] + n(1- p)
n
= p n(n + 1) + n( 1 - p) = p(n + 1) + n(l- p).
n
2
2
This general formula yields some quite reasonable answers. For example, if p = 1
(i.e., the search must be successful), the average number of key comparisons made
by sequential search is (n + 1) /2; i.e., the algorithm will inspect, on average, about
half of the list's elements. If p = 0 (i.e., the search must be unsuccessful), the
average number of key comparisons will be n because the algorithm will inspect
all n elements on all such inputs.
As you can see from this very elementary example, investigation of the
average-case efficiency is considerably more difficult than investigation of the
worst-case and best-case efficiencies. The direct approach for doing this involves
dividing all instances of size n into several classes so that for each instance of the
class the number of times the algorithm's basic operation is executed is the same.
(What were these classes for sequential search?) Then a probability distribution
of inputs is obtained or assumed so that the expected value of the basic operation's
count can be found.
The technical implementation of this plan is rarely easy, however, and probabilistic assumptions underlying it in each particular case are usually difficult
to verify. Given our quest for simplicity, we will mostly quote known results
about average-case efficiency of algorithms under discussion. If you are interested
in derivations of these results, consult such books as [BaaOO], [Sed96], [Knui],
[Knuii], and [Knuiii].
Does one really need the average-case efficiency information? The answer is
unequivocally yes: there are many important algorithms for which the averagecase efficiency is much better than the overly pessimistic worst-case efficiency
would lead us to believe. So, without the average-case analysis, computer scientists
could have missed many important algorithms. Finally, it should be clear from the
preceding discussion that the average-case efficiency cannot be obtained by taking
the average of the worst-case and the best-case efficiencies. Even though this
average does occasionally coincide with the average-case cost, it is not a legitimate
way of performing the average-case analysis.
Yet another type of efficiency is called amortized efficiency. It applies not to
a single run of an algorithm but rather to a sequence of operations performed
on the same data structure. It turns out that in some situations a single operation
can be expensive, but the total time for an entire sequence of n such operations is
always significantly better than the worst-case efficiency of that single operation
50
Fundamentals of the Analysis of Algorithm Efficiency
1I
g
I
multiplied by n. So we cau "amortize" the high cost of such a worst-case occurrence over the entire sequence in a manner similar to the way a business would
amortize the cost of an expensive item over the years of the item's productive life.
This sophisticated approach was discovered by the American computer scientist
Robert Tarjan, who used it, among other applications, in developing an interesting variation of the classic binary search tree (see [Tar87] for a quite readable
nontechnical discussion and [Tar85] for a technical account). We will see an example of the usefulness of amortized efficiency in Section 9.2, when we consider
algorithms for finding unions of disjoint sets.
Recapitulation of the Analysis Framework
Before we leave this section, let us summarize the main points of the framework
outlined above .
.,
.I
Both time and space efficiencies are measured as functions of the algorithm's
input size.
"'
"
Time efficiency is measured by counting the number of times the algorithm's
basic operation is executed. Space efficiency is measured by counting the
number of extra memory units consumed by the algorithm.
The efficiencies of some algorithms may differ significantly for inputs of the
same size. For such algorithms, we need to distinguish between the worst-case,
average-case, and best -case efficiencies .
.,
The framework's primary interest lies in the order of growth of the algorithm's
running time (extra memory units consumed) as its input size goes to infinity.
In the next section, we look at formal means to investigate orders of growth. In
Sections 2.3 and 2.4, we discuss particular methods for investigating nonrecursive
and recursive algorithms, respectively. It is there that you will see how the analysis
framework outlined here can be applied to investigating efficiency of specific
algorithms. You will encounter many more examples throughout the rest of the
book.
-----Exercises
llllliiiiiiilooii!iliiiiiiilliiiiiiiO__,_
1 ---------------
1. For each of the following algorithms, indicate (i) a natural size metric for its
inputs; (ii) its basic operation; (iii) whether the basic operation count can be
different for inputs of the same size:
a. computing the sum of n numbers
b. computing n!
c. finding the largest element in a list of n numbers
d. Euclid's algorithm
IIJ•
2.1
Analysis Framework
51
e. sieve of Eratosthenes
f. pen-and-pencil algorithm for multiplying two 11-digit decimal integers
2. a. Consider the definition-based algorithm for adding two 11-by-11 matrices.
What is its basic operation? How many times is it performed as a function
of the matrix order n? As a function of the total number of elements in the
input matrices?
b. Answer the same questions for the definition-based algorithm for matrix
multiplication.
3. Consider a variation of sequential search that scans a list to return the number
of occurrences of a given search key in the list. Will its efficiency differ from
the efficiency of classic sequential search?
4. a. Glove selection There are 22 gloves in a drawer: 5 pairs of red gloves, 4
pairs of yellow, and 2 pairs of green. You select the gloves in the dark and
can check them only after a selection has been made. What is the smallest
number of gloves you need to select to have at least one matching pair in
the best case? in the worst case? (after [MosOl], 1118)
b. Missing socks Imagine that after washing 5 distinct pairs of socks, you
discover that two socks are missing. Of course, you would like to have
the largest number of complete pairs remaining. Tims, you are left with
4 complete pairs in the best-case scenario and with 3 complete pairs in
the worst case. Assuming that the probability of disappearance for each
of the 10 socks is the same, find the probability of the best-case scenario;
the probability of the worst -case scenario; the number of pairs you should
expect in the average case. (after [Mos01], 1148)
5. a. Prove formula (2.1) for the number of bits in the binary representation of
a positive decimal integer.
b. What would be the analogous formula for the number of decimal digits?
c. Explain why, within the accepted analysis framework, it does not matter
whether we use binary or decimal digits in measuring n 's size.
6. Suggest how any sorting algorithm can be augmented in a way to make the
best-case count of its key comparisons equal to just n - 1 (11 is a list's size,
of course). Do you think it would be a worthwhile addition to any sorting
algorithm?
7. Gaussian elimination, the classic algorithm for solving systems of n linear
equations in 11 unknowns, requires about ~11 3 multiplications, which is the
algorithm's basic operation.
a. How much longer should you expect Gaussian elimination to work on a
system of 1000 equations versus a system of 500 equations?
b. You are considering buying a computer that is 1000 times faster than the
one you currently have. By what factor will the faster computer increase
•
,, ·I··.
:i'•.. I i
52
Fundamentals of the Analysis of Algorithm Efficiency
the sizes of systems solvable in the same amount of time as on the old
computer?
8. For each of the following functions, indicate how much the function's value
will change if its argument is increased fourfold.
b. fo
c. n
d. n 2
e. n 3
f. 2"
a. log2 n
9. Indicate whether the first function of each of the following pairs has a smaller,
same, or larger order of growth (to within a constant multiple) than the second
function.
a. n(n + 1) and 2000n 2
b. 100n 2 and 0.01n3
c. log2 n and In n
e. 2n-l and 211
d. log~ n and log2 n 2
f. (n-l)!andn!
10. Invention of chess According to a well-known legend, the game of chess was
invented many centuries ago in northwestern India by a sage named Shashi.
When he took his invention to his king, the king liked the game so much that
he offered the inventor any reward he wanted. Sashi asked for some grain to
be obtained as follows: just a single grain of wheat was to be placed on the first
square of the chess board, two on the second, four on the third, eight on the
fourth, and so on, until all64 squares had been filled. What would the ultimate
result of this algorithm have been?
Asymptotic Notations and
Basic Efficiency Classes
As pointed out in the previous section, the efficiency analysis framework concentrates on the order of growth of an algorithm's basic operation count as the
principal indicator of the algorithm's efficiency. To compare and rank such orders
of growth, computer scientists use three notations: 0 (big oh), Q (big omega), and
19 (big theta). First, we introduce these notations informally, and then, after several examples, formal definitions are given. In the following discussion, t(n) and
g(n) can be any nonnegative functions defined on the set of natural numbers. In
the context we are interested in, t(n) will be an algorithm's running time (usually
indicated by its basic operation count C(n) ), and g(n) will be some simple function
to compare the cmmt with.
Informal Introduction
Informally, 0 (g(n)) is the set of all functions with a smaller or same order of growth
as g(n) (to within a constant multiple, as n goes to infinity). Thus, to give a few
examples, the following assertions are all true:
1
-n(n- 1) E O(n 2).
2
lI
I
I
l
I
I
J
2.2 Asymptotic Notations and Basic Efficiency Classes
53
Indeed, the first two functions are linear and hence have a smaller order of growth
than g(11) = 11 2, while the last one is quadratic and hence has the same order of
growth as 11 2 On the other hand,
Indeed, the functions 11 3 and 0.0000ln 3 are both cubic and hence have a higher
order of growth than 11 2 ; and so has the fourth-degree polynomial 11 4 + n + 1.
The second notation, O(g(n)), stands for the set of all functions with a larger
or same order of growth as g(11) (to within a constant multiple, as 11 goes to infinity).
For example,
but lOOn+ 5 rfc O(n 2 ).
Finally, EJ(g(11)) is the set of all functions that have the same order of growth
as g(n) (to within a constant multiple, as n goes to infinity). Thus, every quadratic
function a11 2 + bn + c with a > 0 is in C"l(11 2 ), but so are, among infinitely many
others, 11 2 +sin nand 11 2 +log n. (Can you explain why?)
Hopefully, the preceding informal discussion has made you comfortable with
the idea behind the three asymptotic notations. So now come the formal definitions.
a-notation
DEFINITION 1 A function t(n) is said to be in O(g(n)), denoted t(n) E O(g(n)),
if t(11) is bounded above by some constant multiple of g(n) for all large 11, i.e., if
there exist some positive constant c and some nonnegative integer n0 such that
t(n) 00 cg(n)
for all 11 2: 11 0 .
The definition is illustrated in Figure 2.1 where, for the sake of visual clarity, n is
extended to be a real number.
As an example, let us formally prove one of the assertions made in the
introduction: lOOn+ 5 E 0(11 2). Indeed,
lOOn+ 5 S lOOn + n (for all n 2: 5) = lOln S 101n 2
Thus, as values of the constants c and n 0 required by the definition, we can take
101 and 5, respectively.
Note that the definition gives us a lot of freedom in choosing specific values
for constants c and n 0 . For example, we could also reason that
lOOn + 5 S lOOn + 5n (for all 11 2: 1) = l05n
to complete the proof with c = 105 and n 0 = 1.
54
Fundamentals of the Analysis of Algorithm Efficiency
cginl
tin)
doesn't
matter
FIGURE 2.1 Big-oh notation: t(n) E O(g(n))
Q-notation
DEFINITION 2 A function t(n) is said to be in Q (g(n)). denoted t(n) E Q (g(n)),
if t(n) is bounded below by some positive constant multiple of g(n) for all large n,
i.e., if there exist some positive constant c and some nonnegative integer n 0 such
that
t(n) ": cg(n)
for all n ":no.
The definition is illustrated in Figure 2.2.
tin)
cginl
I
doesn't
matter
L-----~---------------------+1!
no
FIGURE 2.2 Big-omega notation: t(n) E Q(g(n))
..•.
2.2
Asymptotic Notations and Basic Efficiency Classes
55
Here is an example of the formal proof that n 3 E Q (n 2 ):
3
2
n 2: n
for all n 2: 0,
i.e., we can select c = 1 and n0 = 0.
8-notation
DEFINITION 3 A function t(n) is said to be in E>(g(n)), denoted t(n) E E>(g(n)),
if t(n) is bounded both above and below by some positive constant multiples of
g(n) for all large n, i.e., if there exist some positive constant c1 and c2 and some
nonnegative integer n 0 such that
c2g(n)
s t(n) s qg(11)
for all 11 2: 11 0 .
The definition is illustrated in Figure 2.3.
For example, let us prove that ~n(n- 1) E 8(11 2). First, we prove the right
inequality (the upper bound):
~11(n- 1) = 2~11 2 - ~n
< ~11 2 for all 11 > 0.
2 -2
-
2
Second, we prove the left inequality (the lower bound):
1
1
11
11
2
2
2 -2
2 2
1
-n(11- 1) = -n 2 - -n > -11 2 - -n-11 (for all n > 2) = -n 2 .
Hence, we can select c2 = ~, c1 =
-
i, 'and n0 = 2.
c1 g(n)
tin)
c2 g(n)
doesn't :
matter
L_----~---------------------+n
no
FIGURE 2.3 Big-theta notation: t(n) E G(g(n))
4
56
Fundamentals of the Analysis of Algorithm Efficiency
Useful Property Involving the Asymptotic Notations
Using the formal definitions o£ the asymptotic notations, we can prove their
general properties (see Problem 7 in Exercises 2.2 for a few simple examples). The
following property, in particular, is useful in analyzing algorithms that comprise
two consecutively executed parts.
THEOREM
If 11(n) E O(g 1(n)) and 12(n) E O(g2(n)), then
l1(n) + l2(n) E O(max{gt(n), g2(n))).
(The analogous assertions are true for the 0 and B notations as well.)
PROOF (As you will see, the proof extends to orders of growth the following
simple fact about four arbitrary real numbers al> bl> a2o and b2 : if a 1 :: b 1 and
a2 :: b2, then a 1 + a 2 :: 2 max{b 1, h2].) Since IJ(n) E O(gJ(n)), there exist some
positive constant c 1 and some nonnegative integer n1 such that
11(n)c:c 1g 1(n)
forallno:n 1.
Similarly, since 12(n) E O(g2(n)),
12(n) C: c2g2(n)
for alln 2: n 2 .
Let us denote c 3 = max{c 1, c2) and consider n 2: max{n 1, n 2] so that we can use
both inequalities. Adding the two inequalities above yields the following:
i](n) + 12(11) :" CJ8J(n) + c2g2(n)
:" c3g1(lr) + c3g2(n) = c3[g1(n) + 82(n)]
I
I
l
C: c32 max{gJ(n), g2(n)).
Hence, 11(n) + 12(n) E O(max{g 1(n), g 2 (n))), with the constants c and n 0 required
by the 0 definition being 2c3 = 2 max{c 1, c2] and max{n 1, n2], respectively.
II
So what does this property imply for an algorithm that comprises two consecutively executed parts? It implies that the algorithm's overall efficiency is determined by the part with a larger order of growth, i.e., its least efficient part:
t1 (n) E O(g1(n)) }
(n) E O(g (n))
12
2
IJ(n) + t2(n) E O(max{gJ(n), gz(n))).
For example, we can check whether an array has identical elements by means of
the following two-part algorithm: first, sort the array by applying some known
sorting algorithm; second, scan the sorted array to check its consecutive elements
for equality. If, for example, a sorting algorithm used in the frrst part makes no
more than ~n(n- 1) comparisons (and hence is in O(n 2)) while the second part
makes no more than n- 1 comparisons (and hence is in O(n)), the efficiency of
the entire algorithm will be in O(max{n 2, n)) = O(n 2).
i
I.
I
II
I
I
2.2
Asymptotic Notations and Basic Efficiency Classes
57
Using limits for Comparing Orders of Growth
Though the formal definitions of 0, Q, and 0J are indispensable for proving their
abstract properties, they are rarely used for comparing the orders of growth of
two specific functions. A much more convenient method for doing so is based on
computing the limit of the ratio of two functions in question. Three principal cases
may arise:
( )
{ 0
lim t__rl_ = c > 0
00
"~ g(n)
oo
implies that t(n) has a smaller order of growth than g(n)
implies that t(n) has the same order of growth as g(n)
implies that t(n) has a larger order of growth than g(n) 3
Note that the first two cases mean that t(n) E O(g(n)), the last two mean that
t(n) E rl(g(n)), and the second case means that t(n) E C"J(g(n)).
The limit -based approach is often more convenient than the one based on
the definitions because it can take advantage of the powerful calculus techniques
developed for computing limits, such as L'Hopital's rule
lim t(n) = lim t'(n)
11--+oo g(n)
11--+oo g'(n)
and Stirling's formula
n!"'
~ ( ~ )" for large values of n.
Here are three examples of using the limit-based approach to comparing
orders of growth of two functions.
EXAMPLE 1 Compare the orders of growth of ~n(n - 1) and n 2 . (This is one of
the examples we used at the beginning of this section to illustrate the definitions.)
ln(n- 1)
lim ~2--..,.-n--+oo
n2
2
1
lim " -n =!lim (1-!)=!.
2 n--+oo n2
2 n--+oo
n
2
Since the limit is equal to a positive constant, the functions have the same order
of growth or, symbolically, in(n- 1) E C"J(n 2).
!!!
EXAMPLE 2 Compare the orders of growth of log2 nand fo. (Unlike Example
1, the answer here is not immediately obvious.)
.
log n
.
(Iog2 n )'
hm - -2= I n n
,
11--+oo
3.
,Jii
11->oo
(,Jit)
lim (Iogz e) ~ = 2 log e lim
n--+oo
_1_
2.j/i
2
n--+oo
Jii = 0.
n
The fourth case, in which such a limit does not exist, rarely happens in the actual practice of analyzing
algorithms. Still, this possibility makes the limit-based approach to comparing orders of growth less
general than the one based on the definitions of 0, ~,and e.
58
Fundamentals of the Analysis of Algorithm Efficiency
Since the limit is equal to zero, log 2 n has a smaller order of growth than fo. (Since
lim Iogb" = 0, we can use the so-called little-oh notation: log2 n E o(fo). Unlike
n---+oo -vn
the big-oh, the little-oh notation is rarely used in analysis of algorithms.)
Ill
EXAMPLE 3 Compare the orders of growth of 11! and 2". (We discussed this issue
informally in the previous section.) Taking advantage of Stirling's formula, we get
lim !!.:_I = lim
/1----:>-00
,I
211
ll--+00
v'2Jrn (~e )" = lim v'21fn-"-"- = lim v'21fn ( !!_ )" = oo.
211
n-HX)
211 ell
11--+00
2e
Thus, though 2" grows very fast, 11! grows still faster. We can write symbolically
that n! E rl(2"); note, however, that while big-omega notation does not preclude
the possibility that n! and 2" have the same order of growth, the limit computed
here certainly does.
Ill
Basic Efficiency Classes
Even though the efficiency analysis framework puts together all the functions
whose orders of growth differ by a constant multiple, there are still infinitely many
such classes. (For example, the exponential functions a" have different orders of
growth for different values of base a.) Therefore, it may come as a surprise that
the time efficiencies of a large number of algorithms fall into only a few classes.
These classes are listed in Table 2.2 in increasing order of their orders of growth,
along with their names and a few comments.
You could raise a concern that classifying algorithms by their asymptotic efficiency would be of little practical use since the values of multiplicative constants
are usually left unspecified. This leaves open a possibility of an algorithm in a
worse efficiency class running faster than an algorithm in a better efficiency class
for inputs of realistic sizes. For example, if the running time of one algorithm is n3
while the running time of the other is 106n2 , the cubic algorithm will outperform
the quadratic algorithm unless n exceeds 10°. A few such anomalies are indeed
known. For example, there exist algorithms for matrix multiplication with a better
asymptotic efficiency than the cubic efficiency of the definition-based algorithm
(see Section 4.5). Because of their much larger multiplicative constants, however,
the value of these more sophisticated algorithms is mostly theoretical.
Fortunately, multiplicative constants usually do not differ that drastically. As
a rule, you should expect an algorithm from a better asymptotic efficiency class
to outperform an algorithm from a worse class even for moderately sized inputs.
This observation is especially true for an algorithm with a better than exponential
running time versus an exponential (or worse) algorithm.
2.2 Asymptotic Notations and Basic Efficiency Classes
59
TABLE 2.2 Basic asymptotic efficiency classes
Class
Name
Comments
1
constant
Short of best-case efficiencies, very few reasonable
examples can be given since an algorithm's running
time typically goes to infinity when its input size grows
infinitely large.
Iogn
logarithmic
Typically, a result of cutting a problem's size by a
constant factor on each iteration of the algorithm (see
Section 5.5). Note that a logarithmic algorithm cannot
take into account all its input (or even a fixed fraction
of it): any algorithm that does so will have at least
linear running time.
n
linear
Algorithms that scan a list of size n (e.g., sequential
search) belong to this class.
n log n
"n-log-n"
Many divide-and-conquer algorithms (see Chapter 4),
including mergesort and quicksort in the average case,
fall into this category.
n2
quadratic
n3
cubic
Typically, characterizes efficiency of algorithms with
two embedded loops (see the next section). Elementary sorting algorithms and certain operations on
n-by-n matrices are standard examples.
JYpically, characterizes efficiency of algorithms with
three embedded loops (see the next section). Several
nontriVial algorithms from linear algebra fall into this
class.
2"
exponential
Typical for algorithms that generate all subsets of an
n-element set. Often, the term "exponential" is used
in a broader sense to include this and larger orders of
growth as well.
n!
factorial
Typical for algorithms that generate all permutations
of ann-element set.
-----Exercises 2 . 2 - - - - - - - - - - - - - - - 1. Use the most appropriate notation among 0. 8, and Q to indicate the time
efficiency class of sequential search (see Section 2.1)
a. in the worst case.
b. in the best case.
c. in the average case.
60
Fundamentals of the Analysis of Algorithm Efficiency
2. Use the informal definitions of 0, El, and Q to determine whether the following assertions are true or false.
a. n(n+1)/2E O(n 3)
c. n(n+l)j2EEl(n 3 )
b. n(n+l)j2E O(n 2 )
d. n(n+1)j2EQ(n)
3. For each of the following functions, indicate the class El(g(n)) the function
belongs to. (Use the simplest g(n) possible in your answers.) Prove your
assertions.
a. (n 2 + 1) 10
c. 2n Jg(n + 2) 2 + (n + 2) 2 Jg!!
e. Llog2 nj
b. Vl0n 2 + 7n + 3
d. zn+l+3n-l
4. a. Table 2.1 contains values of several functions that often arise in analysis of
algorithms. These values certainly suggest that the functions
, 3
n ,2 ", 1
n.
1og n, n, n Iog n, n 2
are listed in increasing order of their order of growth. Do these values
prove this fact with mathematical certainty?
b. Prove that the functions are indeed listed in increasing order of their order
of growth.
5. Order the following functions according to their order of growth (from the
lowest to the highest):
(n- 2)!, Slg(n + 100) 10 , 22", 0.00ln 4 + 3n 3 + 1, Jn 2 n, ..y;J, 3".
6. a. Prove that every polynomial of degree k, p(n) = aknk + ak_ 1nk-l + · · · +
ao, with ak > 0 belongs to El(nk).
b. Prove that exponential functions a" have different orders of growth for
different values of base a > 0.
7. Prove (by using the definitions of the notations involved) or disprove (by
giving a specific counterexample) the following assertions.
a. If t(n) E O(g(n)), then g(n) E Q(l(n)).
b. El(ag(n)) = El(g(n)) where a> 0.
c. El(g(n)) = O(g(n)) n Q(g(n)).
d. For any two nonnegative functions t(n) and g(n) defined on the set of
nonnegative integers, either t(n) E O(g(n)), or t(n) E Q(g(n)), or hath.
8. Prove the section's theorem for
a. Q-notation.
b. 8-notation.
9. We mentioned in this section that one can check whether all elements of an
array are distinct hy a two-part algorithm based on the array's presorting.
2.3
Mathematical Analysis of Nonrecursive Algorithms
61
a. If the presorting is done by an algorithm with the time efficiency in
El(n log n), what will be the time efficiency class of the entire algorithm?
b. If the sorting algorithm used for presorting needs an extra array of size n,
what will be the space efficiency class of the entire algorithm?
You are facing a wall that stretches infinitely in both directions. There is a door in the wall, but you know neither how far away nor in
which direction. You can see the door only when you are right next to it. Design an algorithm that enables you to reach the door by walking at most 0 (n)
steps where n is the (unknown to you) number of steps between your initial
position and the door. [Par95], #652
10. Door in a wall
2.3 Mathematical Analysis of
Nomecursive Algorithms
In this section, we systematically apply the general framework outlined in Section 2.1 to analyzing the time efficiency of nonrecursive algorithms. Let us start
with a very simple example that demonstrates all the principal steps typically taken
in analyzing such algorithms.
EXAMPLE 1 Consider the problem of finding the value of the largest element
in a list of n numbers. For simplicity, we assume that the list is implemented as
an array. The following is a pseudocode of a standard algorithm for solving the
problem.
ALGORITHM MaxE!ernent(A[O .. n -1])
//Determines the value of the largest element in a given array
//Input: An array A[O .. n - 1] of real numbers
//Output: The value of the largest element in A
rnaxva[ +-- A[O]
fori +-lton-1do
if A[i] > rnaxval
maxval +-- A[i]
return maxval
The obvious measure of an input's size here is the number of elements in the
array, i.e., n. The operations that are going to be executed most often are in the
algorithm's for loop. There are two operations in the loop's body: the comparison
A[i] >max val and the assignment max val +-- A[i]. Which of these two operations
should we consider basic? Since the comparison is executed on each repetition
of the loop and the assignment is not, we should consider the comparison to be
the algorithm's basic operation. (Note that the number of comparisons will be the
62
l
Fundamentals of the Analysis of Algorithm Efficiency
!
same for all arrays of size n; therefore, in terms of this metric, there is no need to
distinguish among the worst, average, and best cases here.)
Let us denote C(n) the number of times this comparison is executed and try
to find a formula expressing it as a function of size n. The algorithm makes one
comparison on each execution of the loop, which is repeated for each value of the
loop's variable i within the bounds 1 and n - 1 (inclusively). Therefore, we get the
following sum for C(n):
n-l
C(n) =
I
l
I
I
L 1.
i=l
This is an easy sum to compute because it is nothing else but 1 repeated n - 1
times. Thus,
1
Il-l
C(n) =
L 1=n -lE G(n).
Ill
J
i=l
l
Here is a general plan to follow in analyzing nonrecursive algorithms.
General Plan for Analyzing Time Efficiency of Nonrecursive Algorithms
l,
2.
3.
4.
5.
Decide on a parameter (or parameters) indicating an input's size.
Identify the algorithm's basic operation. (As a rule, it is located in its innermost loop.)
Check whether the number of times the basic operation is executed depends
only on the size of an input. If it also depends on some additional property,
the worst-case, average-case, and, if necessary, best-case efficiencies have to
be investigated separately.
Set up a sum expressing the number of times the algorithm's basic operation
is executed.4
Using standard formulas and rules of sum manipulation, either find a closedform formula for the count or, at the very least, establish its order of growth.
Before proceeding with further examples, you may want to review Appendix
A, which contains a list of summation formulas and rules that are often useful in
analysis of algorithms. In particular, we use especially frequently two basic rules
of sum manipulation
u
u
Lcai=c Lai
i=l
(Rl)
u
i=l
u
u
i=l
i=f
i=l
(R2)
4.
Sometimes, an analysis of a nonrecursive algorithm requires setting up not a sum but a recurrence
relation for the number of times its basic operation is executed. Using recurrence relations is much
more typical for analyzing recursive algorithms (see Section 2.4).
l
!
I
I
I
!I
i
l
2.3
Mathematical Analysis of Nonrecursive Algorithms
63
and two summation formulas
L" 1 = u - l + 1 where l
<( u are some lower and upper integer limits
(Sl)
i=l
~.
L.., z
i=O
~.
n(n + 1)
1 2
2
= L..,
1 = 1 + 2 + · ··+ n =
""-n E El(n ).
i=l
2
(S2)
2
(Note that the formula 'L;':;} 1 = n - 1, which we used in Example 1, is a special
case of formula (Sl) for l = 1 and u = n -1.)
EXAMPLE 2 Consider the element uniqueness problem: check whether all the
elements in a given array are distinct. This problem can be solved by the following
straightforward algorithm.
ALGORITHM
UniqueElements(A[O .. n -1])
//Determines whether all the elements in a given array are distinct
//Input: An array A[O .. n -1]
//Output: Returns "true" if all the elements in A are distinct
II
and "false" otherwise
for i <-- 0 to n - 2 do
for j <-- i + 1 to n - 1 do
if A[i] = A[j] return false
return true
The natural measure of the input's size here is again the number of elements
in the array, i.e., n. Since the innermost loop contains a single operation (the comparison of two elements), we should consider it as the algorithm's basic operation.
Note, however, that the number of element comparisons will depend not only on
n but also on whether there are equal elements in the array and, if there are, which
array positions they occupy. We will limit our investigation to the worst case only.
By definition, the worst ease input is an array for which the number of element
comparisons Cwors 1 (n) is the largest among all arrays of size n. An inspection of
the innermost loop reveals that there are two kinds of worst-case inputs (inputs
for which the algorithm does not exit the loop prematurely): arrays with no equal
elements and arrays in which the last two elements are the only pair of equal
elements. For such inputs, one comparison is made for each repetition of the
innermost loop, i.e., for each value of the loop's variable j between its limits i + 1
and n - 1; and this is repeated for each value of the outer loop, i.e., for each value
of the loop's variable i between its limits 0 and n - 2. Accordingly, we get
rr
,,, ''
,; i!
64
Fundamentals of the Analysis of Algorithm Efficiency
n-2 n-1
Cwm>t(n) =
n-2
n-2
L L 1 = L[(n- 1)- (i + 1) + 1] = L(n -1- i)
i=O j=i+l
=
n-2
n-2
n-2
i=O
i=O
i=O
L (n - 1) - L i = (n - 1) L 1 -
(n _ 2)(n _ 1)
2
= (n _ 1)2 _ (n- 2)(n- 1) = (n- 1)n ""~n2 E G(n2).
2
2
2
We also could have computed the sum I:;'~r~(n - 1- i) faster as follows:
n-2
~
.
~~-Dn
L}n - 1- 1) = (n - 1) + (11 - 2) + · · · + 1 =
.
2
1=0
.I
.
Note that this result was perfectly predictable: in the worst case, the algorithm
needs to compare all n(n- 1)/2 distinct pairs of its n elements.
ll
n-by-n
EXAMPLE 3 Given two
matrices A and B, find the time efficiency of the
definition-based algorithm for computing their product C = AB. By definition, C
is an n-by-n matrix whose elements are computed as the scalar (dot) products of
the rows of matrix A and the columns of matrix B:
A
B
c
col.j
where C[i, j] = A[i, O]B[O, j] + · · · + A[i, k]B[k, j] + · · · + A[i, n -1]B[n -1, j]
for every pair of indices 0 ~ i, j :::=: n - 1.
ALGORITHM
MatrixMultiplication(A[O .. n- 1, O.. n- 1], B[O.. n- 1, O.. n- 1])
//Multiplies two n-by-n matrices by the definition-based algorithm
//Input: Two n-by-n matrices A and B
//Output: Matrix C = AB
fori +-Oton -1do
for j +- 0 to n - 1 do
C[i, j] +- 0.0
for k +- 0 to n - 1 do
C[i, j] +- C[i, j] + A[i, kj * B[k, j]
return C
2.3
Mathematical Analysis of Nonrecursive Algorithms
65
We measure an input's size by matrix order n. In the algorithm's innermost
loop are two arithmetical operations-multiplication and addition-that, in principle, can compete for designation as the algorithm's basic operation. We consider
multiplication as the algorithm's basic operation (see Section 2.1). Note that for
this algorithm, we do not have to choose between these two operations because
on each repetition of the innermost loop, each of the two is executed exactly once.
So by counting one we automatically count the other. Let us set up a sum for the
total number of multiplications M(n) executed by the algorithm. (Since this count
depends only on the size of the input matrices, we do not have to investigate the
worst-case, average-case, and best-case efficiencies separately.)
Obviously, there is just one multiplication executed on each repetition of the
algorithm's innermost loop, which is governed by the variable k ranging from the
lower bound 0 to the upper bound n - 1. Therefore, the number of multiplications
made for every pair of specific values of variables i and j is
n-1
I).
k~O
and the total number of multiplications M(n) is expressed by the following triple
sum:
n-1 n-1 n-1
M(n) L L L 1.
=
i=O j=O k=O
Now we can compute this sum by using formula ( S1) and rule ( R1) (see above).
Starting with the innermost sum L.:~;;;6 1, which is equal ton (why?), we get
n-1 n-1n-1
n-1 n-1
n-1
M(n)= LLL1= LLn= Ln2=n3.
i=O j=O
(This example is simple enough so that we could get this result without all the summation machinations. How? The algorithm computes n 2 elements of the product
matrix. Each of the product's elements is computed as the scalar (dot) product
of an n-element row of the first matrix and an n-element column of the second
matrix, which takes n multiplications. So the total number of multiplications is
n . n 2 = n 3 . It is this kind of reasoning we expected you to employ when answering
this question in Problem 2 of Exercises 2.1.)
If we now want to estimate the running time of the algorithm on a particular
machine, we can do it by the product
T(n) ~ c111 M(n) = c111 n3,
where C111 is the time of one multiplication on the machine in question. We would
get a more accurate estimate if we took into account the time spent on the
additions, too:
r
rlrl
! ,.,
66
Fundamentals of the Analysis of Algorithm Efficiency
where ca is the time of one addition. Note that the estimates differ only by their
multiplicative constants, not by their order of growth.
II!
You should not have the erroneous impression that the plan outlined above
always succeeds in analyzing a nonrecursive algorithm. An irregular change in a
loop's variable, a sum too complicated to analyze, and the difficulties intrinsic to
the average-case analysis are just some of the obstacles that can prove to be insurmountable. These caveats notwithstanding, the plan does work for many simple
nonrecursive algorithms, as you will see throughout the subsequent chapters of
the book.
As a last example, let us consider an algorithm in which the loop's variable
changes in a different manner from that of the previous examples.
EXAMPLE 4 The following algorithm finds the number o£ binary digits in the
binary representation of a positive decimal integer.
ALGORITHM Binary(n)
//Input: A positive decimal integer n
//Output: The number of binary digits in n 's binary representation
count +-1
while n > 1 do
count +---- count + 1
n +- Ln/2J
return count
First, notice that the most frequently executed operation here is not inside the
while loop but rather the comparison n > 1 that determines whether the loop's
body will be executed. Since the number of times the comparison will be executed
is larger than the number of repetitions of the loop's body by exactly 1, the choice
is not that important.
A more significant feature of this example is the fact that the loop's variable
takes on only a few values between its lower and upper limits; therefore we
have to use an alternative way of computing the number of times the loop is
executed. Since the value of n is about halved on each repetition of the loop,
the answer should be about log2 n. The exact formula for the number of times
the comparison 11 > 1 will be executed is actually Llog2 11j + 1-the number of bits
in the binary representation of 11 according to formula (2.1). We could also get
this answer by applying the analysis technique based on recurrence relations; we
discuss this technique in the next section because it is more pertinent to the analysis
II!
of recursive algorithms.
2.3
Mathematical Analysis of Nonrecursive Algorithms
67
-----Exercises 2 . 3 - - - - - - - - - - - - - - 1. Compute the following sums.
a. 1+3+5+7+···+999
b. 2 + 4 + 8 + 16 + ... + 1024
c.
z::;·:i 1
1
f. 2.::}~1 31+
d. z::;·:i i
e. z::;·,..;z i (i + 1)
g. z::;·~l 2.::}~1 ij
h. z::;·~11/i(i + 1)
2. Find the order of growth of the following sums.
2
a.
+1) 2
b.
lg i 2
z::;·,..;tu
I:;',..;i
c. z::;'~ 1 Ci + 1J2'-1
d. z::;•,..;J I:}~1u + n
Use the El(g(n)) notation with the simplest function g(n) possible.
3. The sample variance of n measurements x1, ... , xn can be computed as
"" (
-)2 where i = -"-~i=~1~x_;
""
~"-~i=~1~x-'...,-,-x_
n -1
n
or
I:;'~, xf- CI:7d x;)2 jn
n-1
Find and compare the number of divisions, multiplications, and additions/
subtractions (additions and subtractions are usually bunched together) that
are required for computing the variance according to each of these formulas.
4. Consider the following algorithm.
ALGORITHM
Mystery(n)
//Input: A nonnegative integer n
s of- 0
for i +-- 1 to n do
S<-S+i*i
returnS
a. What does this algorithm compute?
b. What is its basic operation?
c. How many times is the basic operation executed?
d. What is the efficiency class of this algorithm?
e. Suggest an improvement or a better algorithm altogether and indicate its
efficiency class. If you cannot do it, try to prove that in fact it cannot be
done.
r
•
68
Fundamentals of the Analysis of Algorithm Efficiency
5. Consider the following algorithm.
ALGORITHM
Secret(A[O .. n -1])
//Input: An array A[O .. n - 1] of n real numbers
minval +- A[O]; maxval +- A[O]
for i +- 1 to n - 1 do
if A[i] <min val
min val <--- A[i]
if A[i] > maxval
maxval +- A[i]
return maxval - minval
Answer questions (a)-( e) of Problem 4 about this algorithm.
6. Consider the following algorithm.
ALGORITHM
Enigma(A[O .. n -1, O.. n -1])
//Input: A matrix A[O .. n - 1, O.. n - 1] of real numbers
for i +- 0 to n - 2 do
for j +- i + 1 to n - J do
if A[i, j] oft A[j, i]
return false
return true
Answer questions (a)-( e) of Problem 4 about this algorithm.
7. Improve the implementation of the matrix multiplication algorithm (see Example 3) by reducing the number of additions made by the algorithm. What
effect will this change have on the algorithm's efficiency?
8. Determine the asymptotic order of growth for the total number of times all
the doors are toggled in the locker doors puzzle (Problem 11 in Exercises 1.1 ).
9. Prove the formula
~.
L.., 1 = 1
i=l
n(n+l)
+ 2 + · · · + n = --'--c'--'-
2
either by mathematical induction or by following the insight of a 10-year old
schoolboy named Karl Friedrich Gauss (1777-1855), who grew up to become
one of the greatest mathematicians of all times.
10. Consider the following version of an important algorithm that we will study
later in the book.
2.4
Mathematical Analysis of Recursive Algorithms
AlGORITHM
69
GE(A[O.. n -1, O.. nj)
//Input: An n-by-n + 1 matrix A[O .. n - 1, O.. n] of real numbers
for i <- 0 to n - 2 do
for j <- i + 1 to n - 1 do
for k +-- i to n do
A[j, k] <- A[j, k]- A[i, k] * A[j, i] / A[i, i]
a. Find the time efficiency class of this algorithm.
b. What glaring inefficiency does this pseudocode contain and how can it be
eliminated to speed the algorithm up?
11. von Neumann's neighborhood How many one-by-one squares are generated by the algorithm that starts with a single square and on each of its n
iterations adds new squares all round the outside? [Gar99], p. 88. (In the parlance of cellular automata theory, the answer is the number of cells in the
von Newmann neighborhood of range n.) The results for n = 0, 1, and 2 are
illustrated below.
2
Mathematical Analysis of Recursive Algorithms
In this section, we will see how to apply the general framework for analysis
of algorithms to recursive algorithms. We start with an example ofteu used to
introduce novices to the idea of a recursive algorithm.
EXAMPLE 1 Compute the factorial fuuction F(n) = n! for an arbitrary nonnegative integer n. Since
n!=1· ... ·(n-l)·n=(n-l)!·n
forn::o:l
and 0! = 1 by definition, we can compute F(n) = F(n- 1) · n with the following
recursive algorithm.
70
Fundamentals of the Analysis of Algorithm Efficiency
ALGORITHM
F(n)
//Computes n! recursively
//Input: A nonnegative integer n
//Output: The value of n!
if n = 0 return 1
else return F(n - 1) * n
For simplicity, we consider n itself as an indicator of this algorithm's input size
(rather than the number of bits in its binary expansion). The basic operation of the
algorithm is multiplication,5 whose number of executions we denote M(n). Since
the function F(n) is computed according to the formula
for n > 0,
F(n) = F(n- 1) · n
the number of multiplications M (n) needed to compute it must satisfy the equality
M(n) =M(n -1) +
1
to compute
to multiply
F(tJ-·l)
F(n-1) by n
for n > 0.
Indeed, M(n- 1) multiplications are spent to compute F(n- 1), and one more
multiplication is needed to multiply the result by n.
The last equation defines the sequence M (n) that we need to find. This equation defines M(n) not explicitly, i.e., as a function of n, but implicitly as a function
of its value at another point, namely n - 1. Such equations are called recurrence
relations or, for brevity, recurrences. Recurrence relations play an important role
not only in analysis of algorithms b!lt also in some areas of applied mathematics.
They are usually studied in detail in courses on discrete mathematics or discrete
structures; a very brief tutorial on them is provided in Appendix B. Our goal now
is to solve the recurrence relation M (n) = M (n - 1) + 1, i.e., to find an explicit
formula for M(n) in terms of n only.
Note, however, that there is not one but infinitely many sequences that satisfy
this recurrence. (Can you give examples of, say, two of them?) To determine a
solution uniquely, we need an initial condition that tells us the value with which
the sequence starts. We can obtain this value by inspecting the condition that
makes the algorithm stop its recursive calls:
if n = 0 return 1.
This tells us two things. First, since the calls stop when n = 0, the smallest value
of n for which this algorithm is executed and hence M (n) defined is 0. Second, by
inspecting the pseudocode's exiting line, we can see that when n = 0, the algorithm
performs no multiplications. Thus, the initial condition we are after is
5.
Alternatively, we could count the number of times the comparison n = 0 is executed, which is the same
as counting the total number of calls made by the algorithm (sec Problem 2 in Exercises 2.4).
2.4
Mathematical Analysis of Recursive Algorithms
71
M{O) =0.
the calls stop when n = 0 - - - -
' - - - - no multiplications when n = 0
Thus, we succeed in setting up the recurrence relation and initial condition
for the algorithm's number of multiplications M (n):
M(n) = M(n- 1) + 1
for n > 0,
(2.2)
M(O) = 0.
Before we embark on a discussion of how to solve this recurrence, let us
pause to reiterate an important point. We arc dealing here with two recursively
defined functions. The first is the factorial function F(n) itself; it is defined by the
recurrence
F(n) = F(n- 1) · n
for every n > 0,
F(O) = L
The second is the number of multiplications M(n) needed to compute F(n) by the
recursive algorithm whose pseudocode was given at the beginning of the section.
As we just showed, M(n) is defined by recurrence (2.2). And it is recurrence (2.2)
that we need to solve now.
Though it is not difficult to "guess" the solution (what sequence starts with 0
when n = 0 and increases by 1 at each step?), it will be more useful to arrive at it in
a systematic fashion. Among several techniques available for solving recurrence
relations, we use what can be called the method of backward substitutions. The
method's idea (and the reason for the name) is in1mediatcly clear from the way it
applies to solving our particular recurrence:
M(n) = M(n- 1) + 1
= [M(n- 2) + 1] + 1 = M(n- 2) + 2
substitute M(n- 1) = M(n- 2) + 1
substitute M (n - 2) = M (n - 3) + 1
= [M(n- 3) + 1] + 2 = M(n- 3) + 3.
After inspecting the first three lines, we see an emerging pattern, which makes it
possible to predict not only the next line (what would it be?) but also a general
formula for the pattern: M(n) = M(n- i) + i. Strictly speaking, the correctness of
this formula should be proved by mathematical induction, but it is easier to get
the solution as follows and then verify its correctness.
What remains to be done is to take advantage of the initial condition given.
Since it is specified for n = 0, we have to substitute i = n in the pattern's formula
to get the ultimate result of our backward substitutions:
M(n) = M(n- 1) + 1 = · · · = M(n- i) + i = · · · = M(n- n) + n = n.
You should not be disappointed after exerting so much effort to get this
"obvious" answer. The benefits of the method illustrated in this sinlple example
will become clear very soon, when we have to solve more difficult recurrences.
72
Fundamentals of the Analysis of Algorithm Efficiency
Also note that the simple iterative algorithm that accumulates the product of n
consecutive integers requires the same number of multiplications, and it does so
without the overhead of time and space used for maintaining the recursion's stack.
The issue of time efficiency is actually not that important for the problem of
computing n !, however. As we saw in Section 2.1, the function's values get so large
so fast that we can realistically compute its values only for very small n 's. Again, we
use this example just as a simple and convenient vehicle to introduce the standard
approach to analyzing recursive algorithms.
!II
Generalizing our experience with investigating the recursive algorithm for
computing n !, we can now outline a general plan for investigating recursive algorithms.
General Plan for Analyzing Time Efficiency of Recursive Algorithms
1.
2.
3.
4.
5.
Decide on a parameter (or parameters) indicating an input's size.
Identify the algorithm's basic operation.
Check whether the number of times the basic operation is executed can vary
on different inputs of the same size; if it can, the worst-case, average-case, and
best-case efficiencies must be investigated separately.
Set up a recurrence relation, with an appropriate initial condition, for the
number of times the basic operation is executed.
Solve the recurrence or at least ascertain the order of growth of its solution.
EXAMPLE 2 As our next example, we consider another educational workhorse
of recursive algorithms: the Tower of Hanoi puzzle. In this puzzle, we (or mythical
monks, if you do not like to move disks) have n disks of different sizes and three
pegs. Initially, all the disks are on the first peg in order of size, the largest on the
bottom and the smallest on top. The goal is to move all the disks to the third peg,
using the second one as an auxiliary, if necessary. We can move on1y one disk at a
time, and it is forbidden to place a larger disk on top of a smaller one.
The problem has an elegant recursive solution that is illustrated in Figure 2.4.
To move n > 1 disks from peg 1 to peg 3 (with peg 2 as auxiliary), we first move
recursively n - 1 disks from peg 1 to peg 2 (with peg 3 as auxiliary), then move the
largest disk directly from peg 1 to peg 3, and, finally, move recursively n - 1 disks
from peg 2 to peg 3 (using peg 1 as auxiliary). Of course, if n = 1, we can simply
move the single disk directly from the source peg to the destination peg.
Let us apply the general plan to the Tower of Hanoi problem. The number
of disks n is the obvious choice for the input's size indicator, and so is moving
one disk as the algorithm's basic operation. Clearly, the number of moves M(n)
depends on n only, and we get the following recurrence equation for it:
M(n) = M(n- 1) + 1 + M(n- 1)
for n > 1.
With the obvious initial condition M(1) = 1, we have the following recurrence
relation for the number of moves M(n):
2.4
Mathematical Analysis of Recursive Algorithms
73
3
2
FIGURE 2.4 Recursive solution to the Tower of Hanoi puzzle
M(n) = 2M(n- 1) + 1
for 11 > 1,
(2.3)
M(1) = 1.
We solve this recurrence by the same method of backward substitutions:
M(n) = 2M(n- 1) + 1
sub. M(n-1) = 2M(n- 2) + 1
2
= 2[2M (n - 2) + 1] + 1 = 2 M (n - 2) + 2 + 1
2
sub. M (n - 2) = 2M (n - 3) + 1
3
= 2 [2M (n - 3) + 1] + 2 + 1 = 2 M (n - 3) + 22 + 2 + 1.
The pattern of the first three sums on the left suggests that the next one will be
24 M(n- 4) + 23 + 22 + 2 + 1 and, generally, after i substitutions, we get
M(n) = z' M(n- i) + 2i-l + z'- 2 + · · · + 2 + 1 = z' M(n- i) + zi- 1.
Since the initial condition is specified for 11 = 1, which is achieved fori = n - 1, we
get the following formula for the solution to recurrence (2.3):
M(n)=2"- 1M(n-(n-1))+2"- 1 -1
= 2"- 1M(1) + 2"- 1 -1 = 2"- 1 + 2"- 1 -1 = 2" -1.
Thus, we have an exponential algorithm, which will run for an unimaginably
long time even for moderate values of n (see Problem 5 in Exercises 2.4). This is
not due to the fact that this particular algorithm is poor; in fact, it is not difficult
to prove that this is the most efficient algorithm possible for this problem. It is
I
74
Fundamentals of the Analysis of Algorithm Efficiency
--------------- n
n-1
/~n-2
n-2
~n-1
/~n-2
n-2
FIGURE 2.5 Tree of recursive calls made by the recursive algorithm for the Tower of
Hanoi puzzle
the problem's intrinsic difficulty that makes it so computationally hard. Still, this
example makes an important general point:
One should be careful with recursive algorithms because their succinctness may mask their inefficiency.
When a recursive algorithm makes more than a single call to itself, it is
useful for analysis purposes to construct a tree of its recursive calls. In this tree,
nodes correspond to recursive calls, and we can label them with the value of the
parameter (or, more generally, parameters) of the calls. For the Tower of Hanoi
example, the tree is given in Figure 2.5. By counting the number of nodes in the
tree, we can get the total number of talls made by the Tower of Hanoi algorithm:
n-1
C(n) =
L zi
(where lis the level in the tree above)= 2"- 1.
t~o
The number agrees, as it should, with the move count obtained earlier.
!Ill
EXAMPLE 3 As our next example, we investigate a recursive version of the
algorithm discussed at the end of Section 2.3.
ALGORITHM BinRec(11)
//Input: A positive decimal integer 11
//Output: The number of binary digits inn's binary representation
if 11 = 1 return 1
else return BinRec(L11/2J) + 1
Let us set up a recurrence and an initial condition for the number of additions A (11) made by the algorithm. TI1e number of additions made in computing
2.4 Mathematical Analysis of Recursive Algorithms
75
BinRec(Ln/2J) is A(Ln/2J), plus one more addition is made by the algorithm to
increase the returned value by 1. This leads to the recurrence
A(n) = A(Ln/2J) + 1 for 11 > 1.
(2.4)
Since the recursive calls end when n is equal to 1 and there are no additions made
then, the initial condition is
A(1) = 0.
The presence of Ln/2J in the function's argument makes the method of backward substitutions stumble on values of n that are not powers of 2. Therefore, the
standard approach to solving such a recurrence is to solve it only for 11 = 2' and
then take advantage of the theorem called the smoothness rule (see Appendix B),
which claims that under very broad assumptions the order of growth observed for
n = 2' gives a correct answer about the order of growth for all values of n. (Alternatively, after obtaining a solution for powers of 2, we can sometimes fine-tune
this solution to get a formula valid for an arbitrary n.) So let us apply this recipe
to our recurrence, which for n = 2' takes the form
A(2'J = A(2'- 1J + 1
fork> 0,
A(2°) = 0.
Now backward substitutions encounter no problems:
A(2'J = A(2'- 1J + 1
substitute A(2'- 1J = A(2'- 2J + J
2
substitute A(2'- 2J = A(2'- 3 ) + 1
= [A(2'- J + 1] + 1 = A(2'- 2 ) + 2
= (A(2'- 3J + 1] + 2 = A(2'- 3J + 3
Thus, we end up with
A (2') = A (1) + k = k
or, after returning to the original variable n = 2' and hence k = log2 n,
A(n) = log2 n E G(log n).
In fact, we can prove (Problem 6 in Exercises 2.4) that the exact solution for an
arbitrary value ofn is given by just a slightly more refined formula A(n) = Llog2 11j.
!!I
This section provides an introduction to analysis of recursive algorithms.
These techniques will be used throughout the book and expanded further as
necessary. In the next section, we discuss the Fibonacci numbers; their analysis
76
Fundamentals of the Analysis of Algorithm Efficiency
involves more difficult recurrence relations to be solved by a method different
from backward substitutions.
-----Exercises 2 . 4 - - - - - - - - - - - - - - - 1. Solve the following recurrence relations.
a. x(n)=x(n-1)+5 forn>1, x(l)=O
b. x(n) = 3x(n- 1) for n > 1,
x(1) = 4
c. x(n) = x(n- 1) + n for n > 0, x(O) = 0
d. x(n) = x(n/2) + n for n > 1, x(1) = 1 (solve for n = 2k)
e. x(n) = x(n/3) + 1 for n > 1,
x(1) = 1 (solve for n = 3k)
2. Set up and solve a recurrence relation for the number of calls made by F(n),
the recursive algorithm for computing n!.
3. Consider the following recursive algorithm for computing the sum of the first
n cubes: S(n) = 13 + 23 + ... + n 3
ALGORITHM
S(n)
//Input: A positive integer n
//Output: The sum of the first n cubes
if n = 1 return 1
else return S(n - 1) + n * n * n
a. Set up and solve a recurrence relation for the number of times the algorithm's basic operation is executed.
b. How does this algorithm compare with the straightforward nonrecursive
algorithm for computing this sum?
4. Consider the following recursive algorithm.
ALGORITHM
Q(n)
//Input: A positive integer n
if n = 1 return 1
else return Q(n - 1) + 2 * n - 1
a. Set up a recurrence relation for this function's values and solve it to determine what this algorithm computes.
b. Set up a recurrence relation for the number of multiplications made by this
algorithm and solve it.
c. Set up a recurrence relation for the number of additions/subtractions made
by this algorithm and solve it.
;'
'''i,~,,-_, __
2.4
Mathematical Analysis of Recursive Algorithms
77
5. a. Tower of Hanoi In the original version of the Tower of Hanoi puzzle, as
it was published by Edouard Lucas, a French mathematician, in the 1890s,
the world will end after 64 disks have been moved from a mystical Tower
of Brahma. Estimate the number of years it will take if mouks could move
one disk per minute. (Assume that monks do uot eat, sleep, or die.)
b. How many moves are made by the ith largest disk (1:<:: i :s n) in this algorithm?
c. Design a nonrecursive algorithm for the Tower of Hanoi puzzle.
6. a. Prove that the exact number of additions made by the recursive algorithm
BinRec(n) for an arbitrary positive decimal integer n is [log nj.
2
b. Set up a recurrence relation for the number of additions made by the
nonrecursive version of this algorithm (see Section 2.3, Example 4) and
solve it.
7. a. Design a recursive algorithm for computing 2 11 for any nonnegative integer
n that is based on the formula: zn = zn~ 1 + zn- [.
b. Set up a recurrence relation for the number of additions made by the
algorithm and solve it.
c. Draw a tree of recursive calls for this algorithm and count the number of
calls made by the algorithm.
d. Is it a good algorithm for solving this problem?
8. Consider the following recursive algorithm.
ALGORITHM
Minl(A[O .. n -1])
//Input: An array A[O .. n - 1] of real numbers
if n = 1 return A [0]
else temp+- Minl(A[O .. n- 2])
if temp :S A[n-1]returntemp
else return A[n - 1]
a. What does this algorithm compute?
b. Set up a recurrence relation for the algorithm's basic operation count and
solve it.
9. Consider another algorithm for solving the problem of Exercise 8, which
recursively divides an array into two halves: call Min2(A[O .. n - 1]) where
ALGORITHM
Min2(A[l .. r])
if l = r return A [l]
else temp1 +- Min2(A[l .. L(l + r)j2J])
temp2 +- Min2(A[L(l + r)/2J+l..r]l
if templ :S temp2 return templ
else return temp2
78
Fundamentals of the Analysis of Algorithm Efficiency
a. Set up a recurrence relation for the algorithm's basic operation count and
solve it.
b. Which of the algorithms Mini or Min2 is faster? Can you suggest au
algorithm for the problem they solve that would be more efficient than
both of them?
10. The determinant of an n-by-n matrix
all
a21
A=
[
a;li
denoted det A, can be defined as a 11 for n = 1 and, for n > 1, by the recursive
formula
det A=
L"
sjalj det A j•
i=l
where sj is+ 1 if j is odd and -1 if j is even, a1j is the element in row 1 and
column j, and A j is the (n - 1)-by-(n - 1) matrix obtained from matrix A by
deleting its row 1 and column j.
a. Set up a recurrence relation for the number of multiplications made by the
algorithm implementing this recursive definition.
b. Without solving the recurrence, what can you say about the solution's order
of growth as compared to n! ?
11. von Newmann's neighborhood revisited Find the number of cells in the von
Newmann neighborhood of range n (see Problem 11 in Exercises 2.3) by
setting up and solving a recurrence relation.
2.5 Example: Fibonacci Numbers
In this section, we consider the Fibonacci numbers, a famous sequence
0, 1, 1, 2, 3, 5, 8, 13, 21, 34, ...
(2,5)
that can be defined by the simple recurrence
F(n) = F(n -1) + F(n- 2)
for n > 1
(2.6)
and two initial conditions
F(O) = 0, F(1) = 1.
(2.7)
The Fibonacci numbers were introduced by Leonardo Fibonacci in 1202 as a
solution to a problem about the size of a rabbit population. Many more examples
Example: Fibonacci Numbers
2.5
79
of Fibonacci-like numbers have since been discovered in the natural world, and
they have even been used in predicting the prices of stocks and commodities.
There are some interesting applications of the Fibonacci numbers in computer
science as well. For example, worst-case inputs for Euclid's algorithm happen to
be consecutive elements of the Fibonacci sequence. Our discussion goals are quite
limited here, however. First, we find an explicit formula for the nth Fibonacci
number F (n), and then we briefly discuss algorithms for computing it.
Explicit Formula for the nth Fibonacci Number
If we try to apply the method of backward substitutions to solve recurrence (2.6),
we will fail to get an easily discernible pattern. Instead, let us take advantage of a
theorem that describes solutions to a homogeneous second-order linear recurrence
with constant coefficients
ax(n) + bx(n- 1) + cx(n- 2) = 0,
(2.8)
where a, b, and care some fixed real numbers (a -,f 0) called the coefficients of
the recurrence and x (n) is the generic term of an unknown sequence to be found.
According to this theorem-see T11eorem 1 in Appendix B-recurrence (2.8) has
an infinite number of solutions that can be obtained by one of the three formulas.
Which of the three formulas applies to a particular case depends on the number
of real roots of the quadratic equation with the same coefficients as recurrence
(2.8):
ar 2 + br + c = 0.
(2.9)
Quite logically, equation (2.9) is called the characteristic equation for recurrence
(2.8).
Let us apply this theorem to the case of the Fibonacci numbers. To do so,
recurrence (2.6) needs to be rewritten as
F(n)- F(n- 1)- F(n- 2) = 0.
(2.10)
Its characteristic equation is
r
2
-
r -1 = 0,
with the roots
r1,2 =
1± ,/1- 4(-1)
2
l±v'S
2
Since this characteristic equation has two distinct real roots, we have to use the
formula indicated in Case 1 of Theorem 1:
F(n) = 0/
C v'Sr fJ C v'Sr
+2
+
-2
rr...
!·
I
'
80
Fundamentals of the Analysis of Algorithm Efficiency
So far, we have ignored initial conditions (2.7). Now we take advantage of
them to find specific values of parameters a and f3. We do this by substituting 0
and 1-the values of 11 for which the initial conditions are given-into the last
formula and equating the results to 0 and 1 (the values of F(O) and F(1) according
to (2. 7)), respectively:
J5) + f3 e-2J5t = 0
F(1) =a ( l+zJ5Y + f3 e-zJS) = 1.
0
F(O) =a ( 1+2
1
After some standard algebraic simplifications, we get the following system of two
linear equations in two unknowns a and {3:
+
f3
(l+zJS) a + ( l-zJS) f3
a
0
1.
Solving the system (e.g., by substituting f3 = -a into the second eq nation and
solving the equation obtained for a), we get the values a = 1/ vis and f3 = -1/ vis
for the unknowns. Thus,
F(l1) = __.!:_ (1 +vis)" _ _
1 (1- vis)"=__.!:_(¢"_¢,"),
vis
2
.,j5
2
vis
(2.11)
where¢= (1 + vis)/2'" 1.61803 and if,= -11¢'" -0.61803 6 It is hard to believe
that formula (2.11 ), which includes arbitrary integer powers of irrational numbers,
yields nothing else but all the elements of Fibonacci sequence (2.5), but it does!
One of the benefits offormula '(2.11) is that it i=ediately implies that F(n)
grows exponentially (remember Fibonacci's rabbits?), i.e., F(11) E 8(¢"). Tbis
follows from the observation that if, is between -1 and 0, and, hence, ¢" gets
infinitely small as 11 goes to infinity. 1n fact, one can prove that the impact of the
second term)<,¢" on the value of F(11) can be obtained by rounding off the value
of the first term to the nearest integer. 1n other words, for every nonnegative
integer n,
F(11) =
)s¢" rounded to the nearest integer.
•..1•.:
11
(2.12)
J
·.'1:.:
'I
ill
:ll
Algorithms for Computing Fibonacci Numbers
Though the Fibonacci numbers have many fascinating properties, we limit our
discussion to a few remarks about algorithms for computing them. Actually, the
6.
Constant 4> is known as the golden ratio. Since antiquity, it has been considered the most pleasing ratio
of a rectangle's two sides to the human eye and might have been consciously used by ancient architects
and sculptors.
li
,,•t·.l!.
I
~
I
I
1
~~.
~:
2.5
Example: Fibonacci Numbers
81
sequence grows so fast that it is the size of the numbers rather than a time-efficient
method for computing them that should be of primary concern here. Also, for the
sake of simplicity, we consider such operations as additions and multiplications at
unit cost in the algorithms that follow. Since the Fibonacci numbers grow infinitely
large (and grow rapidly), a more detailed analysis than the one offered here is
warranted. These caveats notwithstanding, the algorithms we outline and their
analysis are useful examples for a student of the design and analysis of algorithms.
To begin with, we can use recurrence (2.6) and initial condition (2.7) for the
obvious recursive algorithm for computing F(n).
ALGORITHM
F(n)
//Computes the nth Fibonacci number recursively by using its definition
//Input: A nonnegative integer n
//Output: The nth Fibonacci number
ifn::::; 1 retumn
else return F(n - 1) + F(n - 2)
Before embarking on its formal analysis, can you tell whether this is an efficient algorithm? Well, we need to do a formal analysis anyway. The algorithm's basic operation is clearly addition, so let A (n) be the number of additions performed
by the algorithm in computing F(n). Then the numbers of additions needed for
computing F(n- 1) and F(n- 2) are A(n- 1) and A(n- 2), respectively, and
the algorithm needs one more addition to compute their sum. Thus, we get the
following recurrence for A(n):
A (n) = A (n - 1) + A (n - 2) + 1
for n > 1,
(2.13)
A(O) = 0, A(l) = 0.
The recurrence A(n)- A(n- 1)- A(n- 2) = 1 is quite similar to recurrence
(2.10), but its right-hand side is not equal to zero. Such recurrences are called
inhomogeneous recurrences. There are general techniques for solving inhomogeneous recurrences (see Appendix B or any textbook on discrete mathematics),
but for this particular recurrence, a special trick leads to a faster solution. We can
reduce our inhomogeneous recurrence to a homogeneous one by rewriting it as
[A(n) + 1]- [A(n -1)+ 1]- [A(n- 2) + 1] = 0
and substituting B(n) = A(n) + 1:
B(n)- B(n -1)- B(n -2) =0
B(O) = 1,
B(1) = 1.
Tills homogeneous recurrence can be solved exactly in the same manner as recur-
rence (2.10) was solved to find an explicit formula for F(n). But it can actually be
avoided by noting that B(n) is, in fact, the same recurrence as F(n) except that it
starts with two ones and thus runs one step ahead of F(n). So B(n) = F(n + 1),
rr-
1
!
1
82
'
Fundamentals of the Analysis of Algorithm Efficiency
Fl51
F141~ ~FI31
Fl31/
F121
~Fl11
~FlO)
Fl21
~FlO)
Fl11
Fl21/
~FlO)
~FI11
Fl11
Fl11
FIGURE 2.6 Tree of recursive calls for computing the 5th Fibonacci number by the
definition-based algorithm
and
A(n) = B(n)- 1 = F(n + 1)- 1 =
...!...cq,"+l_ ¢"+1)- 1.
../5
Hence, A(n) E EJ(¢") and, if we measure the size of n by the number of bits
b = Llog2 nj + 1 in its binary representation, the efficiency class will be even worse,
namely doubly exponential. .
The poor efficiency class of the algorithm could be anticipated by the nature of
recurrence (2.13). Indeed, it contains two recursive calls with the sizes of smaller
instances only slightly smaller than size n. (Have you encountered such a situation
before?) We can also see the reason behind the algorithm's inefficiency by looking
at a recursive tree of calls tracing the algorithm's execution. An example of such
a tree for n = 5 is given in Figure 2.6. Note that the same values of the function
are being evaluated again and again, which is clearly extremely inefficient.
We can obtain a much faster algorithm by simply computing the successive
elements of the Fibonacci sequence iteratively, as is done in the following algorithm.
ALGORITHM
Fib(n)
//Computes the nth Fibonacci number iteratively by using its definition
//Input: A nonnegative integer n
//Output: The nth Fibonacci number
F[O] +- 0; F[1] +- 1
for i +- 2 to n do
F[i] +- F[i- 1] + F[i- 2]
return F[n]
This algorithm clearly makes n - 1 additions. Hence, it is linear as a function of n and "only" exponential as a function of the number of bits b in n's
binary representation. Note that using an extra array for storing all the preced-
2.5
Example: Fibonacci Numbers
83
ing elements of the Fibonacci sequence can be avoided: storing just two values is
necessary to accomplish the task (see Problem 6 in Exercises 2.5).
The third alternative for computing the nth Fibonacci numher lies in using
formula (2.12). The efficiency of the algorithm will obviously be determined by
the efficiency of an exponentiation algorithm used for computing¢". If it is done
by simply multiplying ¢ by itself n - 1 times, the algorithm will be in 8 (n) =
e (2b). There are faster algorithms for the exponentiation problem. For example,
we discuss El(log n) = El(b) algorithms for this problem in Chapters 5 and 6.
Note also that special care should be exercised in implementing this approach
to computing the nth Fibonacci number. Since all its intermediate results are
irrational numbers, we would have to make sure that their approximations in the
computer are accurate enough so that the final round-off yields a correct result.
Finally, there exists a El(log n) algorithm for computing the nth Fibonacci
number that manipulates only integers. It is based on the equality
[
F(n- 1)
F(n)
F(n) ]
F(n+1) -
[0 1]"
1 1
for n ::: 1
and an efficient way of computing matrix powers.
------Exercises
1. Find a Web site dedicated to applications of the Fibonacci numbers and
study it.
2. Check by direct substitutions that the function }, (¢" - ¢") indeed satisfies
recurrence (2.6) for every n > 1 and initial conditions (2.7) for n = 0 and 1.
3. The maximum values of the Java primitive types i nt and long are 231 - 1 and
263 - 1, respectively. Find the smallest n for which the nth Fibonacci number
is not going to fit in a memory allocated for
a. the type i nt.
b. the type long.
4. Climbing stairs Find the number of different ways to climb an n-stair staircase if each step is either one or two stairs. (For example, a 3-stair staircase
can be climbed three ways: 1-1-1,1-2, and 2-1.) [Tuc80], p. 112
5. Consider the recursive definition-based algorithm for computing the nth Fibonacci number F (n). Let C (n) and Z (n) be the number of times F (1) and
F(O), respectively, are computed. Prove that
a. C(n) = F(n)
b. Z(n) = F(n- 1).
6. Improve algorithm Fib so that it requires only 8(1) space.
7. Prove the equality
[
F(n- 1)
F(n)
F(n) ]
F(n + 1) -
[0 1]"
1 1
for n 2: 1.
84
Fundamentals of the Analysis of Algorithm Efficiency
8. How many modulo divisious are made by Euclid's algorithm on two consecutive Fibonacci numbers F(n) and F(n- 1) as the algorithm's input?
9. a. Prove Cassini's identity:
F(n + l)F(n -1)- [F(n)f = (-1)''
for n ':: 1.
b. Disappearing square Consider the following paradox, which is based on
Cassini's identity. Take an 8-by-8 chessboard (more generally, any F(n)by-F(n) board divided into [F(n)] 2 squares). Cut it into two trapezoids
and two triangles as shown in the left portion of the figure below. Then
reassemble it as shown in the right portion of the figure. The area of the
left rectangle is 8 x 8 = 64 squares, while the area of the right rectangle is
13 x 5 = 65 squares. Explain the paradox.
1/
---- 1'---
I
1/
1----
1----
1'---
1--------
I
...._
1---- t-...
10. In the language of your choice, implement two algorithms for computing the
last five digits of the nth Fibonacci number that are based on (a) the recursive
definition-based algorithm F(n); (b) the iterative definition-based algorithm
Fih(n). Perform an experiment to find the largest value of n for which your
programs run under 1 minute on your computer.
2.6 Empirical Analysis of Algorithms
In Sections 2.3 and 2.4, we saw how algorithms, both nonrecursive and recursive,
can be analyzed mathematically. Though these techniques can be applied successfully to many simple algorithms, the power of mathematics, even when enhanced
with more advanced techniques (see [Sed96], [Pur85], [Gra94], and [Gre82]), is
far from limitless. In fact, even some seemingly simple algorithms have proved
to be very difficult to analyze with mathematical precision and certainty. As we
pointed out in Section 2.2, this is especially true for average-case analysis.
2.6
Empirical Analysis of Algorithms
85
TIJe principal alternative to the mathematical analysis of an algorithm's efficiency is its empirical analysis. This approach implies the steps spelled out in the
following plan.
General Plan for Empirical Analysis of Algorithm Time Efficiency
1.
2.
3.
4.
Understand the experiment's purpose.
Decide on the efficiency metric M to be measured and the measurement unit
(an operation's count vs. a time unit).
Decide on characteristics of the input sample (its range, size, and so on).
Prepare a program implementing the algorithm (or algorithms) for the experimentation.
5.
6.
7.
Generate a sample of inputs.
Run the algorithm (or algorithms) on the sample's inputs and record the data
observed.
Analyze the data obtained.
Let us discuss these steps one at a time. There are several different goals
one can pursue in analyzing algorithms empirically. They include checking the
accuracy of a theoretical assertion about the algorithm's efficiency, comparing the
efficiency of several algorithms for solving the same problem or different implementations of the same algorithm, developing a hypothesis about the algorithm's
efficiency class, and ascertaining the efficiency of the program implementing the
algorithm on a particular machine. Obviously, an experiment's design should depend on the question the experime~ter seeks to answer.
In particular, the experiment's goal should influence, if not dictate, how the
algorithm's efficiency is to be measured. The fu·st alternative is to insert a counter
(or counters) into a program implementing the algorithm to count the number of
times the algorithm's basic operation is executed. This is usually a straightforward
operation; you should only be mindful of the possibility that the basic operation
is located in several places in the program and that all its executions need to be
accounted for. As straightforward as this task usually is, yon should always test the
modified program to ensure that it works correctly, in terms of both the problem
it solves and the counts it yields.
The second alternative is to time the program implementing the algorithm in
question. The easiest way to do this is to use a system's command, such as the
time conunand in UNIX. Alternatively, we can measure the running time of a
code fragment by asking for the system time right before the fragment's start
Ctstarr) and just after its completion (t finish)) and then computing the difference
between the two (tfinish- tstarr).? InC and C++) you can use the function clock
7.
If the system time is given in units called "ticks," the difference should be divided by a constant
indicating the number of ticks per time unit.
86
Fundamentals of the Analysis of Algorithm Efficiency
for this purpose; in Java, the method cu rrentTi meMill is() in the System class
is available.
It is important to keep several facts in mind, however. First, a system's time
is typically not very accurate, and you might get somewhat different results on
repeated runs of the same program on the same inputs. An obvious remedy is
to make several such measurements and then take their average (or the median)
as the sample's observation point. Second, given the high speed of modern computers, the running time may fail to register at all and be reported as zero. The
standard trick to overcome this obstacle is to run the program in an extra loop
many times, measure the total running time, and then divide it by the number of
the loop's repetitions. Third, on a computer running under a time-sharing system
(such as UNIX), the reported time may include the time spent by the CPU on
other programs, which obviously defeats the purpose of the experiment. Therefore, you should take care to ask the system for the time devoted specifically to
execution of your program. (In UNIX, this time is called the "user time," and it is
automatically provided by the time command.)
Thus, measuring the physical running time has several disadvantages, both
fundamental (dependence on a particular machine being the most important of
them) and technical, not shared by counting the executions of a basic operation.
On the other hand, the physical running time provides very specific information
about an algorithm's performance in a particular computing environment, which
can be of more importance to the experimenter than, say, the algorithm's asymptotic efficiency class. In addition, measuring time spent on different segments of
a program can pinpoint a bottleneck in the program's performance that can be
missed by an abstract deliberation about the algorithm's basic operation. Getting
such data-called profiling-is an important resource in the empirical analysis of
an algorithm's running time; the data in question can usually be obtained from
the system tools available in most computing environments.
Whether you decide to measure the efficiency by basic operation counting
or by time clocking, you will need to decide on a sample of inputs for the experiment. Often, the goal is to use a sample representing a "typical" input; so
the challenge is to understand what a "typical" input is. For some classes of
algorithms-e.g., algorithms for the traveling salesman problem discussed later in
the book-researchers have developed a set of instances they use for benchmarking. But much more often than not, an input sample has to be developed by the
experimenter. Typically, you will have to make decisions about the sample size (it is
sensible to start with a relatively small sample and increase it later if necessary), the
range of input sizes in your sample (typically neither trivially small nor excessively
large), and a procedure for generating inputs in the range chosen. The instance
sizes can either adhere to some pattern (e.g., 1000, 2000, 3000, ... , 10,000 or
500, 1000, 2000, 4000, ... , 128000) or be generated randomly within the range
chosen.
The principal advantage of size changing according to a pattern is that its impact is easier to analyze. For example, if a sample's sizes are generated by doubling,
we can compute the ratios M(2n)jM(n) of the observed metric M (the count or
2.6
Empirical Analysis of Algorithms
87
the time) and see whether the ratios exhibit a behavior typical of algorithms in
one of the basic efficiency classes (see Section 2.2). The major disadvantage of
nonrandom sizes is the possibility that the algorithm under investigation exhibits
atypical behavior on the sample chosen. For example, if all the sizes in a sample
are even and an algorithm under investigation runs much more slowly on odd-size
inputs, the empirical results will be quite misleading.
Another important issue concerning sizes in an experiment's sample is
whether several instances of the same size should be included. If you expect the
observed metric to vary considerably on instances of the same size, it is probably
wise to include several instances for every size in the sample. (TI1ere are welldeveloped methods in statistics to help the experimenter make such decisions;
you will find no shortage of books on this subject.) Of course, if several instances
of the same size are included in the sample, the averages or medians of the observed values for each size should be computed and investigated instead of or in
addition to individual sample points.
Much more often than not, an empirical analysis of an algorithm's efficiency
requires generating random numbers. Even if we decide to use a pattern for input sizes, we typically want instances themselves generated randomly. Generating
random numbers on a digital computer is known to present a difficult problem because, in principle, the problem can be solved only approximately. This is the reason computer scientists prefer to call such numbers pseudorandom. As a practical
matter, the easiest and most natural way of getting such numbers is to take advantage of a random number generator available in computer language libraries.
Typically, its output will be a value of a (pseudo )random variable uniformly distributed in the interval between 0 and 1. If a different (pseudo )random variable
is desired, an appropriate transformation needs to be made. For example, if x is a
continuous random variable uniformly distributed on the interval 0 <( x < 1, the
variable y = l + Lx(r -I)J will be uniformly distributed among the integer values
between integers land r- 1 (I< r).
Alternatively, you can implement one of several known algorithms for generating (pseudo )random numbers. The most widely used and thoroughly studied of
such algorithms is the linear congruential method.
ALGORITHM
Random(n, m, seed, a, b)
//Generates a sequence of 11 pseudorandom numbers according to the linear
//congruential method
//Input: A positive integer n and positive integer parameters m, seed, a, b
//Output: A sequence r 1 , ... , r, of 11 pseudorandom integers uniformly
II
distributed among integer values between 0 and m - 1
//Note: Pseudorandom numbers between 0 and 1 can be obtained
II
by treating the integers generated as digits after the decimal point
r0 +--seed
for i <-- lto n do
ri +----(a*ri-1 +b)modm
I
.
.
"
:1!
.
88
Fundamentals of the Analysis of Algorithm Efficiency
The simplicity of the algorithm's pseudocode is misleading because the devil
lies in the details of choosing the algorithm's parameters. Here is a partial list of
recommendations based on the results of a sophisticated mathematical analysis
(see [Knuii], pp. 184-185, for details): seed may be chosen arbitrarily and is often
set to the current date and time; m should be large and may be conveniently taken
as
zw, where w is the computer's word size; a should be selected as an integer
between O.Olm and 0.99m with no particular pattern in its digits but such that
a mod 8 = 5; the value of b can be chosen as 1.
The empirical data obtained as the result of an experiment need to be recorded
and then presented for an analysis. Data can be presented numerically in a table
or graphically in a scatterplot, that is by points in a Cartesian coordinate system.
It is a good idea to use both these options whenever it is feasible because both
methods have their unique strengths and weaknesses.
The principal advantage of tabulated data lies in the opportunity to manipulate it easily. For example, we can compute the ratios M(n)fg(n) where g(n) is
a candidate to represent the efficiency class of the algorithm in question. If the
algorithm is indeed in E>(g(n)), most likely these ratios will converge to some positive constant as n gets large. (Note that careless novices sometimes assume that
this constant must be 1, which is, of course, incorrect according to the definition
of El(g(n)).) Or we can compute the ratios M(2n)/ M(n) and see how the running
time reacts to doubling of its input size. As we discussed in Section 2.2, such ratios
should change only slightly for logarithmic algorithms and most likely converge
to 2, 4, and 8 for linear, quadratic, and cubic algorithms, respectively-to name
the most obvious and convenient cases.
On the other hand, the form of a scatterplot may also help in ascertaining
the algorithm's probable efficiency class. For a logarithmic algorithm, the scatterplot will have a concave shape (Figure 2.7a); this fact distinguishes it from
all the other basic efficiency classes. For a linear algorithm, the points will tend
to aggregate around a straight line or, more generally, to be contained between
two straight lines (Figure 2.7b ). Scatterplots of functions in E>(n lg n) and E>(n 2)
will have a convex shape (Figure 2.7c), making them difficult to differentiate. A
scatterplot of a cubic algorithm will also have a convex shape, but it will show a
much more rapid increase in the metric's values. An exponential algorithm will
most probably require a logarithmic scale for the vertical axis, in which the values of log" M (n) rather than those of M (n) are plotted. (The commonly used
logarithm base is 2 or 10.) In such a coordinate system, a scatterplot of a truly
exponential algorithm should resemble a linear function because M (n) "" ca" implies 1ogb M (n) "" !ogb c + n !ogb a and vise versa.
One of the possible applications of the empirical analysis is to predict the
algorithm's performance on an instance not included in the experiment sample.
For example, if we observe that the ratios M (n) f g(n) are close to some constant c
for the sample instances, we can approximate M(n) by the product cg(n) for other
instances, too. Though this approach is sensible, it should be used with caution,
especially for values of n outside the sample range. (Mathematicians call such
2.6
count or time
89
Empirical Analysis of Algorithms
count or time
·r
•/
•
• •
•
•
•
•
•
•
•
•
•
•
• •
•
•
n
(b)
FIGURE 2.7 Typical scatterplots: (a) logarithmic; (b) linear; (c) one ofthe convex functions
predictions extrapolation, as opposed to interpolation, which deals with values
within the sample range.) Of course, we cau also try unleashing the standard
techniques of statistical data analysis and prediction. Note, however, that the
majority of such techniques are based on specific probabilistic assumptions that
may or may not be valid for the experimental data in question.
It seems appropriate to end this section hy pointing out the hasic differences hetween mathematical and empirical analyses of algorithms. The principal strength of the mathematical analysis is its independence of specific inputs;
its principal weakness is its limited applicability, especially for investigating the
average-case efficiency. The principal strength of the empirical analysis lies in its
applicability to any algorithm, but its results can depend on the particular sample
of instances and the computer used in the experiment.
J
90
Fundamentals of the Analysis of Algorithm Efficiency
- - - - - Exet·cises 2
1. Consider the following well-known sorting algorithm (we study it more closely
later in the book) with a counter inserted to count the number of key comparisons.
ALGORITHM
SortAnalysis(A[O .. n -1])
//Input: An array A[O .. n -1] of n orderable elements
//Output: The total number of key comparisons made
count +---- 0
for i <-- 1 to n - 1 do
v <-- A[i]
j<--i-1
while j 2: 0 and A[j] > v do
count +-count + 1
A[j + 1] <-- A[i]
j<--j-1
A[j+l]<--v
return count
Is the comparison counter inserted in the right place? If you believe it is, prove
it; if you believe it is not, make an appropriate correction.
2. a. Run the program of Problem 1, with a properly inserted counter (or counters) for the number of key comparisons, on 20 random arrays of sizes 1000,
1500, 2000, 2500, ... , 9000, 9500.
b. Analyze the data obtained to form a hypothesis about the algorithm's
average-case efficiency.
c. Estimate the number of key comparisons one should expect for a randomly
generated array of size 10,000 sorted by the same algorithm.
3. Repeat Problem 2 by measuring the program's running time in milliseconds.
4. Hypothesize a likely efficiency class of an algorithm based on the following
empirical observations of its basic operation's count:
size
1000
2000
3000
4000
5000
6000
7000
8000
9000
10000
count 11,966 24,303 39,992 53,010 67,272 78,692 91,274 113,063 129,799 140,538
5. What scale transformation will make a logarithmic scatterplot look like a
linear one?
6. How can we distinguish a scatterplot for an algorithm in El(lg lg n) from a
scatterplot for an algorithm in 0)(lg n)?
2. 7
Algorithm Visualization
91
7. a. Find empirically the largest number of divisions made by Euclid's algorithm for computing gcd(m, n) for 1::: n :S m :S 100.
b. For each positive integer k, find empirically the smallest pair of integers
1:S 11 :S m s 100 for which Euclid's algorithm needs to make k divisions in
order to find gcd(m, n).
8. The average-case efficiency of Euclid's algorithm on inputs of size 11 can be
measured by the average number of divisions Da,g(n) made by the algorithm
in computing gcd(n, 1), gcd(n, 2), ... , gcd(11, n). For example,
1
Da,g(S) = S(1 + 2 + 3 + 2 + 1) = 1.8.
Produce a scatterplot of Da,,(n) and indicate the algorithm's likely averagecase efficiency class.
9. Run an experiment to ascertain the efficiency class of the sieve of Eratos-
thenes (see Section l.J ).
10. Run a timing experiment for the three algorithms for computing gcd(m, n)
presented in Section l.l.
II
I
I
I
2
Algorithm Visualization
In addition to the mathematical and empirical analyses of algorithms, there is yet
a third way to study algorithms. It is called algorithm visualization and can be
defined as the use of images to convey some useful information about algorithms.
That information can be a visual illustration of an algorithm's operation, of its performance on different kinds of inputs, or of its execution speed versus that of other
algorithms for the same problem. To accomplish this goal, an algorithm visualization uses graphic elements (points, line segments, two- or three-dimensional bars,
and so on) to represent some "interesting events" in the algorithm's operation.
There are two principal variations of algorithm visualization:
'"
lli
static algorithm visualization
dynamic algorithm visualization, also called algorithm animation
Static algorithm visualization shows an algorithm's progress through a series
of still images. Algorithm animation, on the other hand, shows a continuous,
movie-like presentation of an algorithm's operations. Animation is an arguably
more sophisticated option, and it is, of course, much more difficult to implement.
Early efforts in the area of algorithm visualization go back to the 1970s. The
watershed event happened in 1981 with the appearance of a 30-minute color sound
film titled Sorting Out Sorting. The algorithm visualization classic was produced at
the University of Toronto by Ronald Baecker with the assistance of D. Sherman
[Bae81, Bae98]. It contained visualizations of nine well-known sorting algorithms
92
Fundamentals of the Analysis of Algorithm Efficiency
(more than half of them are discussed later in the book) and provided quite a
convincing demonstration of their relative speeds.
The success of Sorting Out Sorting made sorting algorithms a perennial favorite for algorithm animation. Indeed, the sorting problem lends itself quite
naturally to visual presentation via vertical or horizontal bars or sticks of differ-
ent heights or lengths, which are rearranged according to their sizes (Figure 2.8).
This presentation is convenient, however, only for illustrating actions of a typical
sorting algorithm on small inputs. For larger files, Sorting Out Sorting used the
ingenious idea of presenting data by a scatterplot of points on a coordinate plane,
with the first coordinate representing an item's position in the file and the second
one representing the item's value; with such a representation, the process of sort-
ing looks like a transformation of a "random" scatterplot of points into the points
along a frame's diagonal (Figure 2.9). In addition, most sorting algorithms work
by comparing and exchanging two given items at a time-an event that can be
animated relatively easily.
Since the appearance of Sorting Out Sorting, a great number of algorithm animations have been created. They range in scope from one particular algorithm to
a group of algorithms for the same problem (e.g., sorting) or the same application
area (e.g., geometric algorithms) to general-purpose animation systems. The most
widely known of the general-purpose systems include BALSA [Bro84], TANGO
[Sta90], and ZEUS [Bro91]; a· comparative review of their features, along with
those of nine other packages, can be found in [Pri93]. A good general-purpose
animation system should allow a user to not only watch and interact with existing animations of a wide variety of algorithms; it should also provide facilities for
creating new animations. Experience has shown that creating such systems is a
difficult but not impossible task.
The appearance of Java and the World Wiele Web has given a new impetus
to algorithm animation. You are advised to start an exploration with an up-todate site containing a collection of links to sites devoted to algorithm animation.
Since the Web world is notorious for its instability, no specific Web addresses
appear here; a search for the phrase "algorithm auimation" or "algorithm visualization" with a good search engine should do the trick. While you peruse and
evaluate different algorithm animations, you may want to keep in mind the "ten
commandments of algorithm animations." This list of desirable features of an ani-
mation's user interface was suggested by Peter Gloor [Glo98], who was a principal
developer of Animated Algorithms, another well-known algorithm visualization
system:
1. Be consistent.
2. Be interactive.
3. Be clear and concise.
4. Be forgiving to the user.
5, Adapt to the knowledge level of the user.
2.7
Algorithm Visualization
93
FIGURE 2.8 Initial and final screens of a typical visualization of a sorting algorithm using
the bar representation
94
Fundamentals of the Analysis of Algorithm Efficiency
R!ln
II
step
II
Back
II 1'3 II
Extt
l
II
I
..
'
·.·.•.11..
Run
II
step
II
Back
liS II E•~
FIGURE 2.9 Initial and final screens of a typical visualization of a sorting algorithm using
the scatterplot representation
iII
!
I
!!
l,,i
!i
~li
l
ful
------·------
Summary
95
6. Emphasize the visual component.
7. Keep the user interested.
8. Incorporate both symbolic and iconic representations.
9. Include algorithm's analysis (run statistics) and comparisons with other algorithms for the same problem.
10. Include execution history.
There are two principal applications of algorithm visualization: research and
education. The application to education seeks to help students learning algorithms. Potential benefits for researchers are based on expectations that algorithm
visualization may help uncover some unknown features of algorithms. For example, one researcher used a visualization of the recursive Tower of Hanoi algorithm
in which odd- and even-numbered disks were two different colors. He noticed that
two disks of the same color never came in direct contact during the algorithm's execution. This observation helped him in developing a better nonrecursive version
of the classic algorithm.
Although some successes in both education and research applications have
been reported, they are not as impressive as one might expect. Experience has
shown that creating sophisticated software systems is not going to be enough. A
deeper understanding of hum~n perception of images will be required before the
true potential of algorithm visualization is fulfilled.
SUMMARY
There are two kinds of algorithm efficiency: time efficiency and space
" efficiency.
Time efficiency indicates how fast the algorithm runs; space
efficiency deals with the extra space it requires.
An algorithm's time efficiency is principally measured as a function of its
" input
size by counting the number of times its basic operation is executed.
A basic operation is the operation that contributes most toward running
time. Typically, it is the most time-consuming operation in the algorithm's
innermost loop.
.
For some algorithms, the running time may differ considerably for inputs of
the same size, leading to worst-case efficiency, average-case efficiency, and
best-case efficiency.
established framework for analyzing an algorithm's time efficiency is
" The
primarily grounded in the order of growth of the algorithm's running time as
its input size goes to infinity.
notations 0, Q, and 8 are used to indicate and compare the asymptotic
'" The
orders of growth of functions expressing algorithm efficiencies.
.
,
--,(
I
96
Fundamentals of the Analysis of Algorithm Efficiency
m
The efficiencies of a large number of algorithms fall into the following
few classes: constant, logarithmic, linear, "n-log-n," quadratic, cubic, and
exponential.
"
The main tool for analyzing the time efficiency of a nonrecursive algorithm
is to set up a sum expressing the number of executions of its basic operation
and ascertain the sum's order of growth.
"'
The main tool for analyzing the time efficiency of a recursive algorithm is to
set up a recurrence relation expressing the number of executions of its basic
operation and ascertain the solution's order of growth.
il
Succinctness of a recursive algorithm may mask its inefficiency.
m
The Fibonacci numbers are an important sequence of integers in which every
element is equal to the sum of its two immediate predecessors. There are
several algorithms for computing the Fibonacci numbers with drastically
different efficiencies.
"
Empirical analysis of an algorithm is performed by running a program
implementing the algorithm on a sample of inputs and analyzing the data
observed (the basic operation's count or physical running time). This
often involves generating pseudorandom numbers. The applicability to any
algorithm is the principal strength of this approach; the dependence of results
on the particular computer and instance sample is its main weakness.
tm
Algorithm visualization is the use of images to convey useful information
about algorithms. The two principal variations of algorithm visualization are
static algorithm visualization and dynamic algorithm visualization (also called
algorithm animation).
I
I
Brute Force
Science is as far removed from brute force as this sword from a crowbar.
-Edward Lytton (1803-18731, Leila, Book II, Chapter I
Doing a thing well is often a waste of time.
-Robert Byrne, master pool and billiards player and writer
fter introducing the framework and methods for algorithm analysis in the
preceding chapter, we are ready to embark on a discussion of algorithm
design techniques. Each of the next eight chapters is devoted to a particular design
strategy. The subject of this chapter is brute force-the simplest of the design
strategies. lt can be described as follows:
A
Brute force is a straightforward approach to solving a problem, usually
directly based on the problem statement and definitions of the concepts
involved.
The "force" implied by the strategy's definition is that of a computer and
not that of one's intellect. "Just do it!" would be another way to describe the
prescription of the brute-force approach. And often, the brute-force strategy is
indeed the one that is easiest to apply.
As an example, consider the exponentiation problem: compute a" for a given
number a and a nonnegative integer n. Though this problem might seem trivial,
it provides a useful vehicle for illustrating several algorithm design techniques,
including the brute-force approach. (Also note that computing a" mod m for some
large integers is a principal component of a leading encryption algorithm.) By the
definition of exponentiation,
._______,_,
a 11 =a*···* a .
n times
This suggests simply computing a" by multiplying 1 by an times.
97
Brute Force
98
We have already encountered at least two brute-force algorithms in the book:
the consecutive integer checking algorithm for computing gcd(m, n) (Section LJ)
and the definition-based algorithm for matrix multiplication (Section 2.3). Many
other examples are given later in this chapter. (Can you identify a few algorithms
you already know as being based on the brute-force approach?)
Though rarely a source of clever or efficient algorithms, the brute-force approach should not be overlooked as an important algorithm design strategy. First,
unlike some of the other strategies, brute force is applicable to a very wide variety of problems. (In fact, it seems to be the only general approach for which
it is more difficult to point out problems it cannot tackle.) Second, for some important problems (e.g., sorting, searching, matrix multiplication, string matching),
the brute-force approach yields reasonable algorithms of at least some practical
value with no limitation on instance size. TI1ird, the expense of designing a more
efficient algorithm may be unjustifiable if only a few instances of a problem need
to be solved and a brute-force algorithm can solve those instances with accept,
able speed. Fourth, even if too inefficient in general, a brute-force algorithm can
still be useful for solving small-size instances of a problem. Finally, a brute-force
algorithm can serve an important theoretical or educational purpose, e.g., as a
yardstick with which to judge more efficient alternatives for solving a problem.
3.1
Selection Sort and Bubble Sort
In this section, we consider the application of the brute-force approach to the
problem of sorting: given a list of n orderable items (e.g., numbers, characters
from some alphabet, character strings), rearrange them in nondecreasing order.
As we mentioned in Section 1.3, dozens of algorithms have been developed for
solving this very important problem. You might have learned several of them in
the past. If you have, try to forget them for the time being and look at the problem
afresh.
Now, after your mind is unburdened of previous knowledge of sorting algorithms, ask yourself a question: "What would be the most straightforward method
for solving the sorting problem?" Reasonable people may disagree on the answer
to this question. The two algorithms discussed here-selection sort and bubble
sort-seem to be the two prime candidates. The first of these two algorithms is a
better algorithm overall, and it implements the brute-force approach more clearly.
Selection Sort
We start selection sort by scanning the entire given list to find its smallest element
and exchange it with the first element, putting the smallest element in its final
position in the sorted list. Then we scan the list, starting with the second element,
to find the smallest among the last n - 1 elements and exchange it with the second
element, putting the second smallest element in its final position. Generally, on the
3.1
Selection Sort and Bubble Sort
99
ith pass through the list, which we number from 0 ton - 2, the algorithm searches
for the smallest item among the last n - i elements and swaps it with A;:
A0 SA 1 S···sA,._ 1
II
I A,., ... ,A111if1 , .•• ,A,1_ 1
in their final positions
the last n - i elements
After n - 1 passes, the list is sorted.
Here is a pseudocode of this algorithm, which, for simplicity, assumes that the
list is implemented as an array.
ALGORITHM SelectionSort(A[O.. n -1])
//Sorts a given array by selection sort
//Input: An array A[O .. n - 1] of orderable elements
//Output: Array A[O .. n - 1] sorted in ascending order
for i <- 0 to n - 2 do
min+- i
for} <-i+lton-ldo
if A[j] < A[min] min <- j
swap A[i] and A[min]
As an example, the action of the algorithm on the list 89, 45, 68, 90, 29, 34, 17
is illustrated in Figure 3.1.
The analysis of selection sort is straightforward. The input's size is given by
the number of elements n; the algorithm's basic operation is the key comparison
A[j] < A[min]. The number of times it is executed depends only on the array's size
I 89 45
17 I 45
17
17
17
17
17
29
29
29
29
29
68
68
68
34
34
34
34
90
90
90
90
29
29
45
45
45 I 90
45 68
45 68
34
34
17
89
34 89
68 89
68 89
90 89
89 I 90
FIGURE 3.1 Example of sorting with selection sort. Each line corresponds to one
iteration of the algorithm, i.e., a pass through the list tail to the right of
the vertical bar; an element in bold indicates the smallest element found.
Elements to the left of the vertical bar are in their final positions and are
not considered in this and subsequent iterations.
100
Brute Force
and is given by the following sum:
n-2
C(11)=L
u-2
n·-2
11-l
L 1=L[(11-1)-(i+1J+1]=L(11-1-i).
i=O j=i+l
We have already encountered the last sum in analyzing the algorithm of Example 2
in Section 2.3 (so you should be able to compute it now on your own). Whether
you compute this sum by distributing the summation symbol or by immediately
getting the sum of decreasing integers, the answer, of course, must be the same:
n-2 n-1
n-2
C(n) = "L. "L. 1 = "L..-(11- 1- i) = (11 - 1)n .
2
i=O J=i+l
i=O
Thus, sejyc,ti(l!l §81")isa 8(11 2) ~]gori\hm on all inputs. Note, however, that the
number of keyswaps'is ogly~(1))Qr, more precisely, 11 - 1 (one for each repetition
of the i loop). This property distiriguishes selection sort positively from many other
sorting algorithms.
Bubble Sort
Another brute-force application to the sorting problem is to compare adjacent
elements of the list and exchange them if they are out of order. By doing it
repeatedly, we end up "bubbling up" the largest element to the last position on
the list. The next pass bubbles up the second largest element, and so on until,
after n - 1 passes, the list is sorted. Pass i (0 :0 i :0 n - 2) of bubble sort can be
represented by the following diagram:
Ao, ... , A,; B' AJ+l• ... , An-i-11 An-i :S ... :::; An-1
in their final positions
Here is a pseudocode of this algorithm.
ALGORITHM
BubbleSort(A[O.. n - 1])
//Sorts a given array by bubble sort
//Input: An array A[0 .. 11 - 1] o£ orderable elements
//Output: Array A[0 ..11- 1] sorted in ascending order
for i <-- 0 to n - 2 do
for j <-- 0 to 11 - 2 - i do
if A[j + 1] < A[j] swap A[.i] and A[j + 1]
The action of the algorithm on the list 89, 45, 68, 90, 29, 34, 17 is illustrated
as an example in Figure 3.2.
The number of key comparisons for the bubble-sort version given above is
the same for all arrays of size n; it is obtained by a sum that is ahnost identical to
Selection Sort and Bubble Sort
3.1
?
89
45
45
45
45
45
H
45
45
45
45
H
?
45
89
68
68
68
68
?
H
?
68
68
68
68
H
68
68
89
89
89
89
?
H
'
89 <'->
29
29
29
90
90
90 <'->
29
29
29
'
29
29
29
90 <'->
34
34
'
34
29
89 <'-> 34
34
89
17
34
'
?
H
34
34
34
34
90
17
?
H
17
17
17
89
101
17
17
17
17
17
90
90
90
90
90
etc.
FIGURE3.2 First two passes of bubble sort on the list 89, 45, 68, 90, 29, 34, 17. Anew
line is shown after a swap of two elements is done. The elements to the
right of the vertical bar are in their final positions and are not considered in
subsequent iterations of the algorithm.
the sum for selection sort:
n-2 n-2-i
C(n)=L L
11-2
l=L[(n-2-i)-0+1]
i=O )=0
n-2
= L(n -1-i)
(n- l)n E El(n2).
2
The number of key swaps, however, depends on the input. For the worst case of
decreasing arrays, it is the same as the number of key comparisons:
Swm·."(n) = C(n) =
(n- l)n
2
2
E El(n ).
As is often the case with an application of the brute-force strategy, the first
version of an algorithm obtained can often be improved with a modest amount
of effort. Specifically, we can improve the crude version of bubble sort given
above by exploiting the following observation: if a pass through the list makes
no exchanges, the list has been sorted and we can stop the algorithm (Problem 9a
in Exercises 3.1). Though the new version runs faster on some inputs, it is still
in 8(n 2) in the worst and average cases. In fact, even among elementary sorting
methods, bubble sort is an inferior choice, and, if it were not for its catchy name,
you would probably have never heard of it. However, the general lesson you just
learned is important and worth repeating:
n
.··.,
102
Brute Force
A first application of the brute-force approach often results in an
algorithm that can be improved with a modest amount of effort
-----Exercises 3.1 - - - - - - - - - - - - - - - 1. a. Give an example of an algorithm that should not be considered an application of the brute-force approach.
b. Give an example of a problem that cannot be solved by a brute-force
algorithm.
2. a. What is the efficiency of the brute-force algorithm for computing a" as a
function of n? As a function of the number of bits in the binary representation of n?
b. If you are to compute a 11 mod m where a > 1 and n is a large positive integer,
how would you circumvent the problem of a very large magnitude of a"?
3. For each of the algorithms in Problems 4, 5, and 6 of Exercises 2.3, tell whether
or not the algorithm is based on the brute-force approach.
4. a. Design a brute-force algorithm for computing the value of a polynomial
p ( x ) =·· a11 x n + a11 _1x n-l + · · · + a1x + ao
at a given point x 0 and determine its worst-case efficiency class.
b. If the algorithm you designed is in EJ (n 2 ), design a linear algorithm for this
problem.
c. Is it possible to design an algorithm with a better than linear efficiency for
this problem?
5. Sort the list E, X, A, M, P, L, E in alphabetical order by selection sort.
6. Is selection sort stable? (The definition of a stable sorting algorithm was given
in Section 1.3.)
7. Is it possible to implement selection sort for linked lists with the same EJ(n 2)
efficiency as the array version?
8. Sort the list E, X, A, M, P, L, E in alphabetical order by bubble sort.
9. a. Prove that if bubble sort makes no exchanges on its pass through a list, the
list is sorted and the algorithm can be stopped.
b. Write a pseudocode of the method that incorporates this improvement
c. Prove that the worst-case efficiency of the improved version is quadratic.
10. Is bubble sort stable?
@ Alternating disks You have a row of 2n disks of two colors, n dark and n light.
They alternate: dark, light, dark, light, and so on. You want to get all the dark
disks to the right-hand end, and all the light disks to the left-hand end. The
3.2
Sequential Search and Brute-Force String Matching
103
only moves you are allowed to make are those that interchange the positions
of two neighboring disks.
Design an algorithm for solving this puzzle and determine the number of
moves it takes. (Gar99), p. 75
3.2 Sequential Search and
Brute-Force String Matching
We saw in the previous section two applications of the brute-force approach to the
sorting problem. Here we discuss two applications of this strategy to the problem
of searching. The first deals with the canonical problem of searching for an item
of a given value in a given list. The second is different in that it deals with the
string-matching problem.
Sequential Search
We have already encountered a brute-force algorithm for the general searching
problem: it is called sequential search (see Section 2.1). To repeat, the algorithm
simply compares successive elements of a given list with a given search key until
either a match is encountered (successful search) or the list is exhausted without
finding a match (unsuccessful search). A simple extra trick is often employed in
implementing sequential search: if we append the search key to the end of the list,
the search for the key will have to be successful, and therefore we can eliminate
a check for the list's end on each iteration of the algorithm. Here is a pseudocode
for this enhanced version, with its input implemented as an array.
AlGORITHM Sequentia!Search2(A[O.. n ], K)
//Implements sequential search with a search key as a sentinel
//Input: An array A of n elements and a search key K
//Output: The index of the first element in A[O .. n -1) whose value is
II
equal to K or -1 if no such element is found
A[n]<---K
i~O
while A[i] i' K do
i<---i+l
if i < n return i
else return -1
104
Brute Force
Another straightforward improvement can be incorporated in sequential
search if a given list is known to be sorted: searching in such a list can be stopped
as soon as an element greater than or equal to the search key is encountered.
Sequential search provides an excellent illustration of the brute-force approach, with its characteristic strength (simplicity) and weakness (inferior efficiency). The efficiency results obtained in Section 2.1 for the standard version of
sequential search change for the enhanced version only very slightly, so that the
algorithm remains linear in both worst and average cases. We discuss later in the
book several searching algorithms with a better time efficiency.
Brute-Force String Matching
Recall the string-matching problem introduced in Section 1.3: given a string of n
characters called the text and a string of m characters (m :': n) called the pattern,
find a substring of the text that matches the pattern. To put it more precisely, we
want to find i-the index of the leftmost character of the first matching substring
in the text-such that ti =Po .... , ti+J = P.i· .... f;+m-1 = Pm-1:
to
text T
t;
t
Po
Pm-1
pattern P
If matches other than the first one need to be found, a string-matching algorithm
can simply continue working until the entire text is exhausted.
A brute-force algorithm for the string-matching problem is quite obvious:
align the pattern against the first m characters of the text and start matching
the corresponding pairs of characters from left to right until either all m pairs
of the characters match (then the algorithm can stop) or a mismatching pair is
encountered. In the latter case, shift the pattern one position to the right and
resume character comparisons, starting again with the first character of the pattern
and its counterpart in the text. Note that the last position in the text which can still
be a begirming of a matching substring is 11 - m (provided the text's positions are
indexed from 0 ton- 1). Beyond that position, there are not enough characters
to match the entire pattern; hence, the algorithm need not make any comparisons
there.
BruteForceStringMatch(T[O.. n -1]. P[O.. m -1])
//Implements brute-force string matching
//Input: An array T[0 .. 11 - 1] of 11 characters representing a text and
II
an array P[O .. m - 1] of m characters representing a pattern
//Output: The index of the first character in the text that starts a
II
matching substring or -1 if the search is unsuccessful
ALGORITHM
3.2 Sequential Search and Brute-Force String Matching
N
N
0 B 0 D y
0 T
N 0 T
N 0 T
N 0 T
N 0
N
N
T
0
N
0
T I
c E D
105
H I
M
T
0
N
T
0
T
FIGURE 3.3 Example of brute-force string matching. (The pattern's characters that are
compared with their text counterparts are in bold type.}
fori +-Oton-mdo
j+-0
while j < m and P(j] = T[i + j] do
j +-j+1
if j = m return i
return -1
An operation of the algorithm is illustrated in Figure 3.3.
Note that for this example, the algorithm shifts the pattern almost always
after a single character comparison. However, the worst case is much worse: the
algorithm may have to make all m comparisons before shifting the pattern, and
this can happen for each of the n - m + 1 tries. (Problem 6 asks you to give a
specific example of such a situation.) Thus, in the worst case, the algorithm is
in 8(nm). For a typical word search in a natural language text, however, we
should expect that most shifts would happen after very few comparisons (check
the example again). Therefore, the average-case efficiency should be considerably
better than the worst-case efficiency. Indeed it is: for searching in random texts,
it has been shown to be linear, i.e., 8(n + m) = 0l(n). There are several more
sophisticated and more efficient algorithms for string searching. The most widely
known of them-by R. Boyer and J. Moore-is outlined in Section 7.2 along with
its simplification suggested by R. Horspool.
-----Exercises 3 . 2 - - - - - - - - - - - - - - - - 1. Find the number of comparisons made by the sentinel version of sequential
search
a. in the worst case.
b. in the average case if the probability of a successful search is p (0 :S p :S 1).
106
Brute Force
2. As shown in Section 2.1, the average number of key comparisons made by
sequential search (without a sentinel, under standard assumptions about its
inputs) is given by the formula
C,wg(n) =
p(n + 1)
+ n(l- p),
2
where p is the probability of a successful search. Determine, for a flxed n,
the values of p (0 :" p :" 1) for which this formula yields the largest value of
Caog(n) and the smallest value of Caog(n).
3. Gadget testing A firm wants to determine the highest floor of its n-story
headquarters from which a gadget can fall with no impact on the gadget's
functionality. The firm has two identical gadgets to experiment with. Design
an algorithm in the best efficiency class you can to solve this problem.
4. Determine the number of character comparisons that will be made by the
brute-force algorithm in searching for the pattern GANDHI in the text
TH ERE.I S_MORE_TO_LI FE_THAN .INCREAS ING.ITS _SPEED
(Assume that the length of the text-it is 47 characters long-is known before
the search starts.)
5. How many comparisons (both successful and unsuccessful) will be made by
the brute-force algorithm in searching for each of the following patterns in
the binary text of one thousand zeros?
a. 00001
b. 10000
c. 01010
6. Give an example of a text oflength n and a pattern of length m that constitutes
a worst-case input for the brute-force string-matching algorithm. Exactly how
many character comparisons will be made for such input?
7. Write a visualization progtam for the brute-force string-matching algorithm.
8. In solving the string-matching problem, would there be any advantage in
comparing pattern and text characters right-to-left instead of left-to-right?
9. Consider the problem of counting, in a given text, the number of substrings
that start with an A and end with a B. (For example, there are four such
substrings in CABAAXBYA.)
a. Design a brute-force algorithm for this problem and determine its efficiency class.
b. Design a more efflcient algorithm for this problem [Gin04].
A popular diversion in the United States, "word find," asks the
player to find each of a given set of words in a square table fllled with single
letters. A word can read horizontally (left or right), vertically (up or down),
or along a 45 degree diagonal (in any of the four directions), formed by
consecutively adjacent cells of the table; it may wrap around the table's
boundaries, but it must read in the same direction with no zigzagging. The
same cell of the table may be used in different words, but, in a given word,
10. Word find
rI
3.3
Closest-Pair and Convex-Hull Problems by Brute Force
107
the same cell may be used no more than once. Write a computer program for
solving this puzzle.
Write a program based on a version of brute-force pattern
matching for playing Battleship (a classic strategy game) on the computer. The
rules of the game are as follows. There are two opponents in the game (in this
case, a human player and the computer). The game is played on two identical
boards (10-by-10 tables of squares) on which each opponent places his or her
ships, not seen by the opponent. Each player has five ships, each of which
occupies a certain number of squares on the board: a destroyer (2 squares), a
submarine (3 squares), a cruiser (3 squares), a battleship (4 squares), and an
aircraft carrier (5 squares). Each ship is placed either horizontally or vertically,
with no two ships touching each other. The game is played by the opponents
taking turns "shooting" at each other's ships. The result of every shot is
displayed as either a hit or a miss. In case of a hit, the player gets to go again and
keeps playing until missing. The goal is to sink all the opponent's ships before
the opponent succeeds in doing it first. (To sink a ship, all squares occupied
by the ship must be hit.)
11. Battleship game
3.3 Closest-Pair and Convex-Hull Problems
by Brute Force
In this section, we consider a straightforward approach to two well-known problems dealing with a finite set of points in the plane. These problems, aside from
their theoretical interest, arise in two important applied areas: computational geometry and operations research.
Closest-Pair Problem
The closest-pair problem calls for finding two closest points in a set of n points.
For simplicity, we consider the two-dimensional case, although the problem can be
posed for points in higher-dimensional spaces as well. We assume that the points in
question are specified in a standard fashion by their (x. y) Cartesian coordinates
and that the distance between two points P; = (x;. Y;) and Pj = (x j, y j) is the
standard Euclidean distance
d(P;, P;) = J(x;- xj) 2 + (Y;- Yj) 2.
The brute-force approach to solving this problem leads to the following obvious algorithm: compute the distance between each pair of clistinct points and
find a pair with the smallest distance. Of course, we do not want to compute the
distance between the same pair of points twice. To avoid doing so, we consider
only the pairs of points (P;, Pj) for which i < j.
108
Brute Force
ALGORITHM
BruteForceClosestPoints(P)
//Finds two closest points in the plane by brute force
//Input: A list P of n (n ": 2) points P1 = (xJ, y 1), ... , P, = (xn, y,)
//Output: Indices index1 and index2 of the closest pair of points
dmin +- oo
for i <-- 1 to n - 1 do
forj <-i+ltondo
d <-- sqrt ((x1 - x) 2 + (y1 - y 1)2 ) 1/sqrt is the square root function
if d < dmin
dmin +- d; indexl +-- i; index2 +-- j
return index1, index2
The basic operation of the algorithm is computing the Euclidean distance
between two points. In the age of electronic calculators with the square-root
button, one might be led to believe that computing the square root is as simple
an operation as, say, addition or multiplication. Actually, it is not. For starters,
even for most integers, square roots are irrational numbers that therefore can
be found only approximately. Moreover, computing such approximations is not a
trivial matter. But, in fact, computing square roots can be avoided! (Can you think
how?) The trick is to realize that we can simply ignore the square root function
and compare the values (x; - x ) 2 + (y1 - y) 2 themselves. We can do this because
the smaller a number of which we take the square root, the smaller its square root,
or, as mathematicians say, the square root function is strictly increasing.
So, if we replaced<-- sqrt((x;- x) 2 + (Y;- y 1)2 ) by dsqr <-- (x;- x 1P +
(Y; - y 1) 2, the basic operation of the algorithm will be squaring a number. The
number of times it will be executed can be computed as follows:
n-1
C(n) =
n
n-1
L L 2 = 2 l::<n- i)
i=l j=i+l
i=l
= 2[ (n - 1) + (n - 2) + · · · + 1) = (n - 1)n E 8 (n 2 ).
In Chapter 4, we discuss ann log n algorithm for this problem.
Convex-Hull Problem
On to the other problem-computing the convex hull. We start with a definition
of a convex set
DEFINITION A set of points (finite or infinite) in the plane is called convex if
for any two points P and Q in the set, the entire line segment with the endpoints
at P and Q belongs to the set.
3.3 Closest-Pair and Convex-Hull Problems by Brute Force
(a)
109
(b)
FIGURE 3.4 (a) Convex sets. (b) Sets that are not convex.
All the sets depicted in Figure 3.4a are convex, and so are a straight line, a triangle,
a rectangle, and, more generally, any convex polygon/ a circle, and the entire
plane. On the other hand, the sets depicted in Figure 3.4b, any finite set of two or
more distinct points, the boundary of any convex polygon, and a circumference
are examples of sets that are not convex.
Now we are ready for the notion of the convex hull. Intuitively, the convex
hull of a set of n points in the plane is the smallest convex polygon that contains all
of them (either inside or on its boundary). If this formulation does not quite fire up
your enthusiasm, consider the problem as one of barricading n sleeping tigers by
a fence of the shortest length. This interpretation is due to D. Hare! [Har92]; it is
somewhat lively, however, because the fence posts have to be erected right at the
spots where some of the tigers sleep! There is another, much tamer interpretation
of this notion. Imagine that the points in question are represented by nails driven
into a large sheet of plywood representing the plane. Take a rubber band and
stretch it to include all the nails, then let it snap into place. The convex hull is the
area bounded by the snapped rubber band (Figure 3.5).
A formal defmition of the convex hull that is applicable to arbitrary sets,
including sets of points that happen to lie on the same line, follows.
1.
By a triangle, rectangle, and, more generally, a convex polygon, we mean here a region, i.e., the set of
points both inside and on the boundary of lhe shape in question.
110
Brute Force
FIGURE 3.5 Rubber-band interpretation of the convex hull
DEFINITION The convex hull of a set S of points is the smallest convex set
containing S. (The "smallest" requirement means that the convex hull of S must
be a subset of any convex set containing S.)
If S is convex, its convex hull is obviously S itself. If S is a set of two points,
its convex hull is the line segment connecting these points. If Sis a set of three
points not on the same line, its convex hull is the triangle with the vertices at the
three points given; if the three points do lie on the same line, the convex hull is
the line segment with its endpoints at the two points that are farthest apart. For
an example of the convex hull for a larger set, see Figure 3.6.
A study of the examples makes the following theorem an expected result.
THEOREM The convex hull of any setS of n > 2 points (not all on the same line)
is a convex polygon with the vertices at some of the points of S. (If all the points
P,~- - - - - - - - ' l p 6
P, •
FIGURE 3.6 The convex hull for this set of eight points is the convex polygon with
vertices at P1, P5, P6, P7, and P3.
3.3
Closest-Pair and Convex-Hull Problems by Brute Force
111
do lie on the same line, the polygon degenerates to a line segment but still with
the endpoints at two points of S.)
The convex-hull problem is the problem of constructing the convex hull for a
given set S of n points. To solve it, we need to find the points that will serve as the
vertices of the polygon in question. Mathematicians call vertices of such a polygon
"extreme points." By definition, an extreme point of a convex set is a point of this
set that is not a middle point of any line segment with endpoints in the set. For
example, the extreme points of a triangle are its three vertices, the extreme points
of a circle are all the points of its circumference, and the extreme points of the
convex hull of the set of eight points in Figure 3.6 are P~o P5, ? 6, P7, and P3.
Extreme points have several special properties other points of a convex set
do not have. One of them is exploited by a very important algorithm called
the simplex method (Section 10.1). This algorithm solves linear programming
problems, problems of finding a minimum or a maximum of a linear function of
n variables subject to linear constraints (see Problem 10 in Exercises 3.3 for an
example and Sections 6.6 and 10.1 for a general discussion). Here, however, we
are interested in extreme points because their identification solves the convex-hull
problem. Actually, to solve this problem completely, we need to know a bit more
than just which of n points of a given set are extreme points of the set's convex
hull: we need to know which pairs of points need to be connected to form the
boundary of the convex hull. Note that this issue can also be addressed by listing
the extreme points in a clockwise or a counterclockwise order.
So how can we solve the convex-hull problem in a brute-force manner? If you
do not see an immediate plan for a frontal attack, do not be dismayed: the convexhull problem is one with no obvious algorithmic solution. Nevertheless, there is a
simple but inefficient algorithm that is based on the following observation about
line segments making up a boundary of the convex hull: a line segment connecting
two points P, and Pj of a set of n points is a part of its convex hull's boundary if and
only if all the other points of the set lie on the same side of the straight line through
these two points.2 (Verify this property for the set in Figure 3.6.) Repeating this
test for every pair of points yields a list of line segments that make up the convex
hull's boundary.
A few elementary facts from analytical geometry are needed to implement
this algorithm. First, the straight line through two points (xJ> y 1), (x 2 , y 2 ) in the
coordinate plane can be defined by the equation
ax+by=c,
where a= Y2- y 1, b = x1 - xz, c = X1Y2 - y,xz.
2.
For the sake of simplicity, we assume here that no three points of a given set lie on the same line. A
modification needed for the general case is left for the exercises.
112
Brute Force
Second, such a line divides the plane into two half-planes: for all the points in
one of them, ax+ by> c, while for all the points in the other, ax+ by< c. (For
the points on the line itself, of course, ax+ by= c.) Thus, to check whether certain
points lie on the same side of the line, we can simply check whether the expression
ax + by - c has the same sign at each of these points. We leave the implementation
details as an exercise.
What is the time efficiency of this algorithm? It is in O(n 3): for each of
n(n- 1)/2 pairs of distinct points, we may need to find the sign of ax+ by- c
for each of the other n - 2 points. There are much more efficient algorithms for
this important problem, and we discuss one of them later in the book.
1. Can you design a faster algorithm than the one based on the brute-force
strategy to solve the closest-pair problem for n points x 1, ••• , x, on the real
line?
2. Let x 1 < x2 < · · · < x11 be real numbers representing coordinates of n villages
located along a straight road. A post office needs to be built in one of these
villages.
a. Design an efficient algorithm to find the post office location minimizing
the average distance between the villages and the post office.
b. Design an efficient algorithm to fmd the post office location minimizing
the maximum distance from a village to the post office.
3. a. There are several alternative ways to define a distance between two points
P, = (xJ, Ytl and P2 = (x2> Y2l in the Cartesian plane. In particular, the
so-called Manhattan distance is defined as
Prove that dM satisfies the following axioms which every distance function
must satisfy:
i. dM(P,, P2 ) 2: 0 for any two points P1 and P2o and dM(P1 , P2 ) = 0 if and
only if P1 = P2 ;
ii. dM(PJ, P2 ) = dM(Pz, Pt);
iii. dM(PJ, Pz) S dM(P,, P3) + dM(P3, Pz) for any P,, P2, and P3.
b. Sketch all the points in the x, y coordinate plane whose Manhattan distance
to the origin (0,0) is equal to 1. Do the same for the Euclidean distance.
c. True or false: A solution to the closest -pair problem does not depend on
which of the two metrics-dE (Euclidean) or dM (Manhattan)-is used?
4. Odd pie fight
There arc n 2: 3 people positioned on a field (Euclidean plane)
so that each has a unique nearest neighbor. Each person has a cream pie. At
3.3
Closest-Pair and Convex-Hull Problems by Brute Force
113
a signal, everybody hurles his or her pie at the nearest neighbor. Assuming
that n is odd and that nobody can miss his or her target, true or false: There
always remains at least one person not hit by a pie [Car79]?
5. The closest-pair problem can be posed on the k-dimensional space in which the
Euclidean distance between two points P' = (x~, ... , x~) and P" = (x~, ... , x;)
is defined as
k
d(P', P") =
2.
""(x'
- x")
~ .Y
s
s=l
What will be the efficiency class of the brute-force algorithm for the kdimensional closest-pair problem?
6. Find the convex hulls of the following sets and identify their extreme points
(if they have any).
a. a line segment
b. a square
c. the boundary of a square
d. a straight line
7. Design a linear-time algorithm to determine two extreme points of the convex
hull of a set of n > 1 points in the plane.
8. What modification needs to be made in the brute-force algorithm for the
convex-hull problem to handle more than two points on the same straight
line?
9. Write a program implementing the brute-force algorithm for the convex-hull
problem.
10. Consider the following small instance of the linear programming problem:
maximize
subject to
3x + 5y
x + y :':' 4
X+ 3y :':' 6
X": 0, y ": 0.
a. Sketch, in the Cartesian plane, the problem's feasible region, defined as
the set of points satisfying all the problem's constraints.
b. Identify the region's extreme points.
c. Solve the optimization problem given by using the following theorem: a
linear programming problem with a nonempty bounded feasible region
always has a solution, which can be found at one of the extreme points of
its feasible region.
!!
i'
114
Brute Force
3.4 Exhaustive Search
Many important problems require finding an element with a special property in a
domain that grows exponentially (or faster) with an instance size. Typically, such
problems arise in situations that involve-explicitly or implicitly-<;ombinatorial
objects such as permutations, combinations, and subsets of a given set. Many such
problems are optimization problems: they ask to find an element that maximizes
or minimizes some desired characteristic such as a path's length or an assignment's
cost.
Exhaustive search is simply a brute-force approach to combinatorial problems. It suggests generating each and every element of the problem's domain,
selecting those of them that satisfy all the constraints, and then finding a desired
element (e.g., the one that optimizes some objective function). Note that though
the idea of exhaustive search is quite straightforward, its implementation typically requires an algorithm for generating certain combinatorial objects. We delay
a discussion of such algorithms until Chapter 5 and assume here that they exist.
We illustrate exhaustive search by applying it to three important problems: the
traveling salesman problem, the knapsack problem, and the assignment problem.
Traveling Salesman Problem
The traveling salesman problem (TSP) has been intriguing researchers for the
last 150 years by its seemingly simple formulation, important applications, and
interesting connections to other combinatorial problems. In layman's terms, the
problem asks to find the shortest tour through a given set of n cities that visits each
city exactly once before returning to the city where it started. The problem can be
conveniently modeled by a weighted graph, with the graph's vertices representing
the cities and the edge weights specifying the distances. Then the problem can be
stated as the problem of finding the shortest Hamiltonian circuit of the graph. (A
Hamiltonian circuit is defined as a cycle that passes through all the vertices of the
graph exactly once. lt is named after the Irish mathematician Sir William Rowan
Hamilton (1805-1865), who became interested in such cycles as an application of
his algebraic discoveries.)
It is easy to see that a Hamiltonian circuit can be also defined as a sequence of
n + 1 adjacent vertices vio• vi 1 , . . . , Vi 1, vio• where the first vertex of the sequence
is the same as the last one while all the other n - 1 vertices are distinct. Further,
we can assume, with no loss of generality, that all circuits start and end at one
particular vertex (they are cycles after all, are they not?). Thus, we can get all
the tours by generating all the permutations of n - 1 intermediate cities, compute
the tour lengths, and find the shortest among them. Figure 3.7 presents a small
instance of the problem and its solution by this method.
An inspection of Figure 3.7 reveals tltree pairs of tours that differ only by their
direction. Hence, we could cut the number of vertex permutations by half. We
could, for example, choose any two intermediate vertices, say, B and C, and then
11
_
I
!
3.4
115
Exhaustive Search
2
a)----'=----{ b
3
5
8
7
c )-------{d
Tour
a -> b -> c -> d -> a
I = 2 + 8 + 1 + 7 = 18
optimal
a -> b --> d -> c -> a
a -> c -> b -> d -> a
a -> c -> d --> b -> a
I = 5 + 8 + 3 + 7 = 23
optimal
a -> d -> b -> c -> a
a -> d -> c -> b -> a
FIGURE 3.7 Solution to a small instance of the traveling salesman problem by exhaustive
search
consider only permutations in which B precedes C. (This trick implicitly defines
a tour's direction.)
This improvement cannot brighten the efficiency picture much, however. The
total number of permutations needed will still be (n- 1)!/2, which makes the
exhaustive-search approach impractical for all but very small values of n. On the
other hand, if you always see your glass as half-full, you can claim that cutting
the work by half is nothing to sneeze at, even if you solve a small instance of the
problem, especially by hand. Also note that had we not limited our investigation
to the circuits starting at the same vertex, the number of permutations would have
been even larger, by a factor of n.
Knapsack Problem
Here is another well-known problem in algorithmics. Given n items of known
weights w1, ... , w" and values v1, ... , v, and a knapsack of capacity W, find the
most valuable subset of the items that fit into the knapsack. If you do not like the
idea of putting yourself in the shoes of a thief who wants to steal the most valuable
116
Brute Force
loot that fits into his knapsack, think about a transport plane that has to deliver
the most valuable set of items to a remote location without exceeding the plane's
capacity. Figure 3.8a presents a small instance of the knapsack problem.
The exhaustive-search approach to this problem leads to generating all the
subsets of the set of n items given, computing the total weight of each subset to
identify feasible subsets (i.e., the ones with the total weight not exceeding the
knapsack's capacity), and finding a subset of the largest value among them. As an
example, the solution to the instance of Figure 3.8a is given in Figure 3.8b. Since
the number of subsets of ann-element set is 2", the exhaustive search leads to a
rl (2") algorithm no matter how efficiently individual subsets are generated.
Thus, for both the traveling salesman and knapsack problems, exhaustive
search leads to algorithms that are extremely inefficient on every input. In
fact, these two problems are the best-known examples of so-called NP-hard
problems. No polynomial-time algorithm is known for any NP-hard problem.
Moreover, most computer scientists believe that such algorithms do not exist, although this very important conjecture has never been proven. More sophisticated
approaches-backtracking and branch-and-bound (see Sections 12.1 and 12.2)enable us to solve some but not all instances of these (and similar) problems in
less than exponential time. Alternatively, we can use one of many approximation
algorithms, such as those described in Section 12.3.
Assignment Problem
In our third example of a problem that can be solved by exhaustive search, there
are n people who need to be assigned to execute n jobs, one person per job. (That
is, each person is assigned to exactly one job and each job is assigned to exactly
one person.) The cost that would accrue if the ith person is assigned to the jthjob
is a known quantity C[i, .i] for each pair i, j = 1, 2, ... , n. The problem is to find
an assignment with the mimimum total cost.
A small instance of this problem follows, with the table entries representing
the assignment costs C[i, j]:
Person 1
Person 2
Person 3
Person 4
Job 1
Job 2
Job 3
Job4
9
6
5
7
2
4
7
3
1
9
8
7
8
6
8
4
It is easy to see that an instance of the assignment problem is completely
specified by its cost matrix C. In terms of this matrix, the problem calls for a
tI
f
3.4
(:
Exhaustive Search
117
I!
I
t
10
w, =7
v, ~ $42
~
w3 = 4
w4 = 5
2
v, ~ $40
v4 ~ $25
item 2
item 3
item4
Subset
Total weigbt
Total value
item 1
knapsack
(a)
0
0
$0
(1)
7
$42
(2)
3
$12
(3)
4
$40
(4)
5
$25
(1, 2)
10
$36
(1, 3)
11
(1, 4)
12
not feasible
not feasible
(2, 3)
7
$52
(2, 4)
8
$37
{3,4}
9
$65
{1, 2, 3)
14
(1, 2, 4)
15
not feasible
not feasible
not feasible
not feasible
not feasible
{1, 3, 4)
16
{2, 3, 4)
12
{1, 2, 3, 4)
19
(b)
FIGURE 3.8 (a) Instance of the knapsack problem. (b) Its solution by exhaustive search.
(The information about the optimal selection is in bold.)
118
Brute Force
selection of one element in each row of the matrix so that all selected elements
are in different columns and the total sum of the selected elements is the smallest
possible. Note that no obvious strategy for finding a solution works here. For
example, we cannot select the smallest element in each row because the smallest
elements may happen to be in tbe same column. In fact, the smallest element in
the entire matrix need not be a component of an optimal solution. Thus, opting
for the exhaustive search may appear as an unavoidable evil.
We can describe feasible solutions to the assignment problem as 11-tuples
(h. ... , .i,) in which the ith component, i = 1, ... , n, indicates the column ofthe
element selected in the ith row (i.e., the job number assigned to the ith person).
For example, for the cost matrix above, (2, 3, 4, 1) indicates a feasible assignment
of Person 1 to Job 2, Person 2 to Job 3, Person 3 to Job 4, and Person 4 to
Job 1. The requirements of the assignment problem imply that there is a one-toone correspondence between feasible assignments and permutations of the first
11 integers. Therefore, the exhaustive-search approach to the assignment problem
would require generating all the permutations of integers 1, 2, ... , n, computing
the total cost of each assignment by summing up the corresponding elements of
the cost matrix, and finally selecting the one with the smallest sum. A few first
iterations of applying this algorithm to the instance given above are shown in
Figure 3.9; you are asked to complete it in the exercises.
Since the number of permutations to be considered for the general case of the
assignment problem is 11 !, exhaustive search is impractical for all but very small
instances of the problem. Fortunately, there is a much more efficient algorithm for
this problem called the Hungarian method after the Hungarian mathematicians
Konig and Egervary whose work underlies the method (see, e.g., [Kol95]).
This is good news: the fact that a problem's domain grows exponentially (or
faster) does not necessarily imply that there can be no efficient algorithm for
solving it. In fact, we present several other examples of such problems later in the
book. However, such examples are more of an exception from the rule. More often
than not, there are no known polynomial-time algorithms for problems whose
domain grows exponentially with instance size (provided we want to solve them
exactly). And, as mentioned, such algorithms quite possibly do not exist.
C= [
~
2
4
8
6
7
3
9
n
<1. 2. 3. 4>
<1' 2. 4. 3>
<1. 3, 2, 4>
<1, 3, 4, 2>
<1,4, 2, 3>
<1,4,3, 2>
COSt=9+4+ 1 +4= 18
cost=9+4+8+9=30
cost = 9 + 3 + 8 + 4 = 24
cost = 9 + 3 + 8 + 6 = 26
cost=9+7+8+9=33
cost = 9 + 7 + 1 + 6 = 23
i
!
etc.
FIGURE 3.9 First few iterations of solving a small instance of the assignment problem
by exhaustive search
!,
1
I,
l~
l
I
I
3.4 Exhaustive Search
119
-----Exercises
1. a. Assuming that each tour can be generated in constant time, what will be
the efficiency class of the exhaustive-search algorithm outlined in the text
for the traveling salesman problem?
b. If this algorithm is programmed on a computer that makes 1 billion additions per second, estimate the maximum number of cities for which the
problem can be solved in
i. one hour.
ii. 24-hours.
iii. one year.
iv. one century.
2. Outline an exhaustive-search algorithm for the Hamiltonian circuit problem.
3. Outline an algorithm to determine whether a counected graph represented
by its adjacency matrix has a Eulerian circuit. What is the efficiency class of
your algorithm?
4. Complete the application of exhaustive search to the instance of the assignment problem started in the text.
5. Give an example of the assignment problem whose optimal solution does not
include the smallest element of its cost matrix.
6. Consider the partition problem: given n positive integers, partition them into
two disjoint subsets with the same sum of their elements. (Of course, the problem does not always have a solution.) Design an exhaustive-search algorithm
for this problem. Try to minimize the number of subsets the algorithm needs
to generate.
7. Consider the clique problem: given a graph G and a positive integer k, determine whether the graph contains a clique of size k, i.e., a complete subgraph
of k vertices. Design an exhaustive-search algorithm for this problem.
8. Explain how exhaustive search can be applied to the sorting problem and
determine the efficiency class of such an algorithm.
9. Magic squares A magic square of order n is an arrangement of the numbers
from 1 to n2 in an n-by-n matrix, with each number occurring exactly once, so
that each row, each column, and each main diagonal has the same sum.
a. Prove that if a magic square of order n exists, the sum in question must be
equal to n(n 2 + 1)/2.
b. Design an exhaustive-search algorithm for generating all magic squares of
order n.
120
Brute Force
c. Go to the Internet or yonr library and find a better algorithm for generating
magic squares.
d. Implement the two algorithms-the exhaustive search and the one you
have found-and run an experiment to determine the largest value of n
for which each of the algorithms is able to find a magic square of order n
in less than one minute of your computer's time.
10. Famous alphametic A puzzle in which the digits in a correct mathematical
expression, such as a sum, are replaced by letters is called a cryptarithm; if,
in addition, the puzzle's words make sense, it is said to be an alphametic.
The most well-known alphametic was published by renowned British puzzlist
H. E. Dudeney (1857-1930):
SEND
+MORE
MONEY
Two conditions are assumed: first, the correspndence between letters and
decimal digits is one-to-one, i.e., each letter represents one digit only and different letters represent different digits; second, the digit zero does not appear
as the left -most digit in any of the numbers. To solve an alphametic means
to find which digit each letter represents. Note that a solution's uniqueness
cannot be assumed and has to be verified by the solver.
a. Write a program for solving cryptarithms by exhaustive search. Assume
that a given cryptarithm is a sum of two words.
b. Solve Dudeney's puzzle the way it was expected to be solved when it was
first published in 1924.
SUMMARY
"'
Brute force is a straightforward approach to solving a problem, usually directly
based on the problem statement and definitions of the concepts involved.
"'
The principal strengths of the brute-force approach are wide applicability and
simplicity; its principal weakness is the subpar efficiency of most brute-force
algorithms.
"
A first application of the brute-force approach often results in an algorithm
that can be improved with a modest amount of effort.
"
The following noted algorithms can be considered as examples of the bruteforce approach:
f.·,,.
-----------------------
< ··r·-
Summary
121
definition-based algorithm for matrix multiplication;
selection sort;
sequential search;
straightforward string-matching algorithm.
"'
Exhaustive search is a brute-force approach to combinatorial problems. It
suggests generating each and every combinatorial object of the problem,
selecting those of them that satisfy all the constraints, and then finding a
desired object.
"
The traveling salesman problem, the knapsack problem, and the assignment
problem are typical examples of problems that can be solved, at least
theoretically, by exhaustive-search algorithms.
"
Exhaustive search is impractical for all but very small instances of problems
it can be applied to.
'
~!"""'"
m<
i"')
Divide-and-Conquer
Whatever man prays for, he prays for a miracle. Every prayer reduces itself
to this~Great God, grant that twice two be not four.
-Ivan Turgenev {1818-1883), Russian novelist and short-story writer
D
ivide-and-conquer is probably the best-known general algorithm design
technique. Though its fame may have something to do with its catchy name, it
is well deserved: quite a few very efficient algorithms are specific implementations
of this general strategy. Divide-and-conquer algorithms work according to the
following general plan:
1.
2.
3.
A problem's instance is divided into several smaller instances of the same
problem, ideally of about the same size.
The smaller instances are solved (typically recursively, though sometimes a
different algorithm is employed when instances become small enough).
If necessary, the solutions obtained for the smaller instances are combined to
get a solution to the original instance.
The divide-and-conquer technique is diagrammed in Figure 4.1, which depicts the
case of dividing a problem into two smaller subproblems, by far the most widely
occurring case (at least for divide-and-conquer algorithms designed to be executed
on a single-processor computer).
As an example, let us consider the problem of computing the sum of n numbers
a 0 , ... , an-1· If n > 1, we can divide the problem into two instances of the same
problem: to compute the sum of the first Ln/2J numbers and to compute the sum
of the remaining fn /21 numbers. (Of course, if n = 1, we simply return a 0 as the
answer.) Once each of these two sums is computed (by applying the same method,
i.e., recursively), we can add their values to get the sum in question:
ao + · · · + an-1 = (ao + · · · + aln/2J-l) + (aln/2J + · · · + an-ll·
Is this an efficient way to compute the sum of n numbers? A moment of
reflection (why could it be more efficient than the brute-force summation?), a
small example of summing, say, four numbers by this algorithm, a formal analysis
123
124
Divide-and-Conquer
problem
of size n
subproblem 1
of size n/2
subproblem 2
of size n/2
solution to
subproblem 1
solution to
subproblem 2
solution to
the original problem
FIGURE 4.1 Divide-and-conquer technique (typical case)
(which follows), and common sense (we do not compute sums this way, do we?)
all lead to a negative answer to this question.
Thus, not every divide-and-conquer algorithm is necessarily more efficient
than even a brute-force solution. But often our prayers to the Goddess of
Algorithmics-see the chapter's epigraph-are answered, and the time spent on
executing the divide-and-conquer plan turns out to be smaller than solving a problem by a different method. In fact, the divide-and-conquer approach yields some of
the most important and efficient algorithms in computer science. We discuss a few
classic examples of such algorithms in this chapter. Though we consider only sequential algorithms here, it is worth keeping in miud that the divide-and-conquer
technique is ideally suited for parallel computations, in which each subproblem
can be solved simultaneously by its own processor.
As mentioned above, in the most typical case of divide-and-conquer, a problem's instance of size n is divided into two instances of size n/2. More generally,
an instance of size n can be divided into b instances of size n/b, with a of them
needing to be solved. (Here, a and b are constants; a 2: 1 and b > 1. ). Assuming
4.1
Mergesort
125
that size n is a power of b, to simplify our analysis, we get the following recurrence
for the running time T(n):
T(n) = aT(11jb) + f(n),
(4.1)
where f(n) is a function that accounts for the time spent on dividing the problem
into smaller ones and on combining their solutions. (For the summation example,
a = b = 2 and f (n) = 1.) Recurrence ( 4.1) is called the geneml divide-and-conquer
recurrence. Obviously, the order of growth of its solution T(n) depends on the
values of the constants a and b and the order of growth of the function .f (n). The
efficiency analysis of many divide-and-conquer algorithms is greatly simplified by
the following theorem (see Appendix B).
THEOREM (Master Theorem)
tion (4.1), then
If .f(n) E 8(11d) with d :> 0 in recurrence equa-
8(11d)
8(11d log n)
T(11) E
{
8(n 10g" ")
if a< hd
if a= h"
if a > bd
(Analogous results hold for the 0 and Q notations, too.)
For example, the recurrence equation for the number of additions A (11) made
by the divide-and-conquer sum-computation algorithm (see above) on inputs of
size n = 2k is
A(11) = 2A(nj2)
+ 1.
Thus, for this example, a= 2, b = 2, and d = 0; hence, since a> bd,
A(11) E 8(n 10g"") = 8(11log2 2 ) = G(n).
Note that we were able to find the solution's efficiency class without going through
the drudgery of solving the recurrence. But, of course, this approach can only establish a solution's order of growth to within an unknown multiplicative constant,
while solving a recurrence equation with a specific initial condition yields an exact
answer (at least for 11 's that are powers of b).
4.1
Mergesort
Mergesort is a perfect example of a successful application of the divide-andconquer technique. It sorts a given array A[O .. n - 1] by dividing it into two halves
A[D .. L11/2J -1] and A[lnf2j .. n -1], sorting each of them recursively, and then
merging the two smaller sorted arrays into a single sorted one.
126
Divide-and-Conquer
ALGORITHM
Mergesort(A[O .. n - 1])
//Sorts array A[O .. n - 1] by recursive mergesort
//Input: An array A[O .. n - 1] of orderable elements
//Output: Array A[O .. n - 1] sorted in nondecreasing order
ifn > 1
copy A[O .. Ln/2J -1]to B[O .. Ln/2J -1]
copy A[Ln/2J .. n -1] to C[o .. rn/21 -1]
Mergesort(B[O .. Ln/2J - 1])
Mergesort(C[O .. rn/21 -1])
Merge(B, C, A)
The merging of two sorted arrays can be done as follows. Two pointers (array
indices) are initialized to point to the first elements of the arrays being merged.
The elements pointed to are compared, and the smaller of them is added to a new
array being constructed; after that, the index of the smaller element is incremented
to point to its immediate successor in the array it was copied from. This operation
is repeated until one of the two given arrays is exhausted, and then the remaining
elements of the other array are copied to the end of the new array.
ALGORITHM
Merge(B[O .. p- 1], C[O .. q -1], A[O.. p + q -1])
//Merges two sorted arrays into one sorted array
//Input: Arrays B[O .. p -1] and C[O .. q -1] both sorted
//Output: Sorted array A[O .. p + q -1] of the elements of Band C
i +--0; j +--0; k +--0
whilei <pandj <q do
if B[i] :S C[j]
A[k] +- B[i]; i +- i + 1
else A[k] +- C[j]; j +- j + 1
k+-k+1
ifi = p
copy C[j .. q -1] to A[k .. p +q -1]
else copy B[i..p -1] to A[k .. p + q -1]
The operation of the algorithm on the list 8, 3, 2, 9, 7, 1, 5, 4 is illustrated in
Figure 4.2.
How efficient is mergesort? Assuming for simplicity that n is a power of 2, the
recurrence relation for the number of key comparisons C(n) is
C(n) = 2C(n/2) + C,,.,,(n)
for n > 1,
C(l) = 0.
Let us analyze C,,.g,(n), the number of key comparisons performed during the
merging stage. At each step, exactly one comparison is made, after which the total
number of elements in the two arrays still needed to be processed is reduced
4.1
Mergesort
127
8 3 2 9 7 1 5 4
12345789
FIGURE 4.2 Example of mergesort operation
by one. In the worst case, neither of the two arrays becomes empty before the
other one contains just one element (e.g., smaller elements may come from the
alternating arrays). Therefore, for the worst case, Cmerge(n) = n- 1, and we have
the recurrence
1 ; 1
Cwom(n) = 2Cwm-, 1 (nj2) + n- 1 for 11 > 1,
Hence, according to the Master Theorem, Cwm-,Jn) E 8(11 log n) (why?). In fact,
it is easy to find the exact solution to the worst-case recurrence for n = 2k:
cwa>.>l(n) = 11 logz n- n
+ 1.
The number of key comparisons made by mergesort in the worst case comes
very close to the theoretical minimum 1 that any general comparison-based sorting
algorithm can have. The principal shortcoming of mergesort is the linear amount
1.
As we shall see in Section 11.2, this theoretical minimum is rtog 2 n !l ~ rn log2 n - 1.44n 1.
Divide-and-Conquer
128
of extra storage the algorithm requires. Though merging can be done in place,
the resulting algorithm is quite complicated and, since it has a significantly larger
multiplicative constant, the in-place mergesort is of theoretical interest only.
-----Exercises 4.1 - - - - - - - - - - - - - - - 1. \;JWrite a pseudocode for a divide-and-conquer algorithm for finding a position of the l!'.!~~t..~ement in an array of n numbers.
b. What will be your algorithm's output for arrays with several elements of
the largest value?
c. Set up and solve a recurrence relation for the number of key comparisons
made by your algorithm.
d. How does this algorithm compare with the brute-force algorithm for this
problem?
2. a. Write a pseudocode for a divide-and-conquer algorithm for finding values
of both the largest and smallest elements in an array of n numbers.
b. Set up and solve (for n = 2k) a recurrence relation for the number of key
comparisons made by your algorithm.
c. How does this algorithm compare with the brute-force algorithm for this
problem?
3. a. Write a pseudocode for a divide-and-conquer algorithm for the exponentiation problem of computing a" where a > 0 and n is a positive integer.
b. Set up and solve a recurrence relation for the number of multiplications
made by this algorithm.
c. How does this algorithm compare with the brute-force algorithm for this
problem?
4. We mentioned in Chapter 2, that logarithm bases are irrelevant in most
contexts arising in the analysis of an algorithm's efficiency class. Is this true
.. for both assertions of the Master Theorem that include logarithms?
,
5. Find the order of growth for solutions of the following recurrences.
•
a. T(n) = 4T(nf2) + n, T(1) = 1
b. T(n) = 4T(n/2) + n 2 , T(1) = 1
'I
c. T(n) = 4T(nf2) + n 3 , T(l) = 1
6. Apply mergesort to sort the list E, X, A, M, P, L, E in alphabetical order.
7. Is mergesort a stable sorting algorithm?
)S._Ja. Solve the recurrence relation for the number of key comparisons made by
mergesort in the worst case. (You may assume that n = 2k.)
'\
\ ':
\
I.
\
\ j\
\\
~
r
I
I
4.2 Quicksort
129
b. Set up a recurrence relation for the number of key comparisons made by
mergesort on best-case inputs and solve it for n = 2k.
c. Set up a recurrence relation for the number of key moves made by the
version of mergesort given in Section 4.1. Does taking the number of key
moves into account change the algorithm's efficiency class?
9. Let A[O .. n - 1] be an array of n distinct real numbers. A pair (A[i], A[j]) is said
to be an inversion if these numbers are out of order, i.e., i < j but A[i] > A[j].
Design an 0 (n log n) algorithm for counting the number of inversions.
10. One can implement mergesort without a recursion by starting with merging
adjacent elements of a given array, then merging sorted pairs, and so on. Implement this bottom-up version of mergesort in the language of your choice.
A tromino is an L-shaped tile formed by 1-by-1 adjacent
squares. The problem is to cover any 2"-by-2" chessboard with one missing
square (anywhere on the board) with trominos. Trominos should cover all the
squares except the missing one with no overlaps.
11. Tromino puzzle
I
it
Design a divide-and-conquer algorithm for this problem.
4.2 Quicksort
j
I
Quicksort is another important sorting algorithm that is based on the divide-andconquer approach. Unlike mergesort, which divides its input's elements according
to their position in the array, quicksort divides them according to their value.
Specifically, it rearranges elements of a given array A(O .. n - 1] to achieve its
partition, a situation where all the elements before some position s are smaller
than or equal to A[s] and all the elements after positions are greater than or equal
toA[s]:
A[O] ... A[s -1] A[s] A[s + 1] ... A[n -1]
all are ~A(s]
all arc ::::A[s]
r
130
Divide-and-Conquer
Obviously, after a partition has been achieved, A[s] will be in its final position in
the sorted array, and we can continue sorting the two subarrays of the elements
preceding and following A[s] independently (e.g., by the same method).
Quicksort(A[Lr])
//Sorts a subarray by quicksort
//Input: A subarray A[Lr] of A[O,n -1], defmed by its left and right indices
II
l and r
//Output: Subarray A[l .. r] sorted in nondecreasing order
if l < r
s +-Partition(A[l .. r]) lis is a split position
Quicksort(A[l .. s- 1])
Quicksort(A[s + l..r])
ALGORITHM
A partition of A[O .. n- 1] and, more generally, ofits subarray A[l .. r] (0::: l <
r ::: n - 1) can be achieved by the following algorithm. First, we select an element
with respect to whose value we are going to divide the subarray. Because of its
guiding role, we call this element the pivot. There are several different strategies
for selecting a pivot; we will return to this issue when we analyze the algorithm's
efficiency. For now, we use the simplest strategy of selecting the subarray's first
element: p = A[l].
There are also several alternative procedures for rearranging elements to
achieve a partition. Here we use an efficient method based on two scans of the
subarray: one is left-to-right and the other right-to-left, each comparing the subarray's elements with the pivot. The left-to-right scan, denoted below by index i,
starts with the second element. Since we want elements smaller than the pivot to
be in the first part of the subarray, this scan skips over elements that are smaller
than the pivot and stops on encountering the first element greater than or equal
to the pivot. The right-to-left scan, denoted below by index j, starts with the last
element of the subarray. Since we want elements larger than the pivot to be in
the second part of the sub array, this scan skips over elements that are larger than
the pivot and stops on encountering the first element smaller than or equal to the
pivot.
After both scans stop, three situations may arise, depending on whether or not
the scanning indices have crossed. If scanning indices i and j have not crossed, i.e.,
i < j, we simply exchange A[i] and A[j] and resume the scans by incrementing i
and decrementing j, respectively:
i~
all are ,; p
I~pI
f
all are 2 p
f
4.2 Quicksort
131
If the scanning indices have crossed over, i.e., i > j, we will have partitioned
the array after exchanging the pivot with A[J]:
Finally, if the scanning indices stop while pointing to the same element, i.e.,
i = j, the value they are pointing to must be equal top (why?). Thus, we have the
array partitioned, with the split positions = i = j:
----i=j--
I p I all are s p 1- pI all are ~ p
We can combine the last case with the case of crossed-over indices (i > j) by
exchanging the pivot with A[J] whenever i 2" j.
Here is a pseudocode implementing this partitioning procedure.
ALGORITHM Partition(A[l .. r])
//Partitions a subarray by using its first element as a pivot
//Input: A subarray A[l..r] of A[O .. n - 1], defined by its left and right
II
indices land r (! < r)
//Output: A partition of A[l..r ], with the split position returned as
II
this ftmction's value
p <-A[l]
i +-l;
j +- r + 1
repeat
repeat i <-- i + 1 until A[i]O> p
repeat j <-- j - 1 until A[j] <( p
swap(A[i], A[j])
until i 2" j
swap(A[i], A[j]) //undo last swap when i 2" j
swap(A[!], A[j])
return j
Note that index i can go out of the subarray bounds in this pseudocode. Rather
than checking for this possibility every time index i is incremented, we can append
to array A[O .. n- 1] a "sentinel" that would prevent index i from advancing beyond
position n. The more sophisticated method of pivot selection mentioned at the end
of the section makes such a sentinel unnecessary.
An example of sorting an array by quicksort is given in Figure 4.3.
\
-r
132
Divide-and-Conquer
2
0
3
4
5
6
7
8
2
4
7
8
2
4
j
i
5
3
9
5
3
9
i
5
5
5
i
3
3
7
j
9
7
9
7
8
9
7
8
8
9
9
7
7
i
j
8
9
7
i
j
7
9
8
2
i
j
4
8
2
j
4
i
2
4
3
j
j
5
4
4
2
3
3
i
j
2
3
4
2
3
2
5
i
2
2
4
i
j
1
3
j
i
1
2
3
3
4
4
4
jj
3
j
3
4
i
4
4
8
8
7
7
j
i
7
8
9
(b)
9
9
(a)
FIGURE 4.3 Example of Quicksort operation. (a) The array's transformations with pivots
shown in bold. (b) The tree of recursive calls to Quicksort with input values
1 and r of subarray bounds and split positions of a partition obtained.
We start our discussion of quicksort's efficiency by noting that the number of
key comparisons made before a partition is achieved is n + 1 if the scanning indices
cross over, n if they coincide (why?). If all the splits happ~ninthe middle of correspondingsubarrays, we will have thelbesfcase:fheiiiimber of key comparisons
inth€ best case will satisfy-the recurrence
4.2 Quicksort
133
According to the Master Theorem, cb"t(n) E 8(n logz n); solving it exactly for
2k yields cb"t(n) = n iogz n.
In the worst case, all the splits will be skewed to the extreme: one of the two
subarrays will be empty, while the size of the other will be just one less than the
size of the subarray being partitioned. This unfortunate situation will happen, in
particular, for increasing arrays, i.e., for inputs for which the problem is already
solved! Indeed, if A[O .. n- 1] is a strictly increasing array and we use A(O] as the
pivot, the left-to-right scan will stop on A[1] while the right-to-left scan will go all
the way to reach A(O], indicating the split at position 0:
11 =
j-- ----i
IA[OJI A[lll
I A[n-lll
So, after making n + 1 comparisons to get to this partition and exchanging the
pivot A(O] with itself, the algorithm will find itself with the strictly increasing array
A(1 .. n - 1] to sort. This sorting of strictly increasing arrays of diminishing sizes will
continue until the last one A[n- 2 .. n -1] has been processed. The total number
of key comparisons made will be equal to
Cwm·,(n)=(n+1)+n+···+3=
.
(n + l)(n + 2)
2
2
-3E8(n ).
Thus, the question about the utility of quicksort comes to its average-case behavior. Let Caag(n) be the average number of key comparisons made by quicksort
on a randomly ordered array of size n. Assuming that the partition split can happen in each positions (0::: s::: n- 1) with the same probability ljn, we get the
following recurrence relation
1 11-1
Caag(n) =- l)(n + 1) + Caag(s) + Cav 8 (n -1- s)]
for n > 1,
n s=O
Ca, 8 (0) = 0, Caag(1) = 0.
Though solving this recurrence is easier tban one might expect, it is still much
trickier than the worst- and best-case analyses, and we willleave it for the exercises.
Its solution turns out to be
Caag(n) "'2n Inn"' 1.38n log2 n.
Thus, on the average, quicksort makes only 38% more comparisons than in the
best case. Moreover, its innermost loop is so efficient that it runs faster than
mergesort (and heapsort, another n log n algorithm that we discuss in Chapter
6) on randomly ordered arrays, justifying the name given to the algorithm by its
inventor, the prominent British computer scientist C.A.R. Hoare.2
2.
The young Hoare invented his algorithm while trying to sort words of a Russian dictionary for a
machine translation project from Russian to English. Says Hoare, "My first thought on how to do
r
'
.
.
.
'
134
Divide-and-Conquer
Given the importance of quicksort, there have been persistent efforts over the
years to refine the basic algorithm. Among several improvements discovered by
researchers are: better pivot selection methods (such as the median-of-three partitioning that uses as a pivot the median of the leftmost, rightmost, and the middle
element of the array); switching to a simpler sort on smaller subfiles; and recursion elimination (so-called nonrecursive quicksort). According toR. Sedgewick
[Sed98], the world's leading expert on quicksort, these improvements in combination can cut the running time of the algorithm by 20%-25%.
We should also point out that the idea of partitioning can be useful in applications other than sorting. In particular, it underlines a fast algorithm for the
important selection problem discussed in Section 5.6.
-----Exercises 4.2 - - - - - - - - - - - - - - - 1. Apply quicksort to sort the list
E, X, A, M, P, L, E
in alphabetical order. Draw the tree of the recursive calls made.
2. For the partitioning procedure outlined in Section 4.2:
a. Prove that if the scanning indices stop while pointing to the same element,
i.e., i = j, the value they are pointing to must be equal top.
b. Prove that when the scanning indices stop, j cannot point to an element
more than one position to the left of the one pointed to by i.
c. Why is it worth stopping the scans after encountering an element equal to
the pivot?
3. Is quicksort a stable sorting algorithm?
4. Give an example of an array of n elements for which the sentinel mentioned
in the text is actually needed. What should be its value? Also explain why a
single sentinel suffices for any input.
5. For the version of quicksort given in the text:
a. Are arrays made up of all equal elements the worst-case input, the bestcase input, or neither?
b. Are strictly decreasing arrays the worst-case input, the best-case input, or
neither?
6. a. For quicksort with the median-of-three pivot selection, are increasing arrays the worst-case input, the best-case input, or neither?
this was bubblesort and, by an amazing stroke of luck, my second thought was Quicksort." It is hard to
disagree with his overall assessment: "I have been very lucky. What a wonderful way to start a career
in Computing, by discovering a new sorting algorithm!" [Hoa96]
4.3
Binary Search
135
b. Answer the same question for decreasing arrays.
7. Solve the average-case recurrence for quicksort.
8. Design an algorithm to rearrange elements of a given array of n real numbers so that all its negative elements precede all its positive elements. Your
algorithm should be both time- and space-efficient.
(9) The Dutch flag problem is to rearrange any array of characters R, W, and
B (red, white, and blue are the colors of the Dutch national flag) so that all
'···
the R's come first, the W's come next, and the B's come last. Design a linear
in-place algorithm for this problem.
10. Implement quicksort in the language of your choice. Run your program on
a sample of inputs to verify the theoretical assertions about the algorithm's
efficiency.
r[j)Nuts and bolts You are given a collection of n bolts of different widths and
n corresponding nuts. You are allowed to try a nut and bolt together, from
which you can determine whether the nut is larger than the bolt, smaller than
the bolt, or matches the bolt exactly. However, there is no way to compare
two nuts together or two bolts together. The problem is to match each bolt to
its nut. Design an algorithm for this problem with average-case efficiency in
G(n log n). [Raw91], p. 293
4.3 Binary Search
Binary search is a remarkably efficient algorithm for searching in a sorted array. It
works by comparing a search key K with the array's middle element A[m].lfthey
match, the algorithm stops; otherwise, the same operation is repeated recursively
for the first half of the array if K < A[m], and for the second half if K > A[m]:
K
t
A[O] ... A[m- 1] A[m] A[m + 1] ... A[n -1].
search here if
K<A[m]
search here if
K>A[m]
As an example, let us apply binary search to searching for K = 70 in the array
136
Divide-and-Conquer
The iterations of the algorithm are given in the following table:
index
value
0 1 2 3 4 5 6 7 8 9 10 11 12
3 II4 127131 139 l42lss 17o 1741 s1lss 193 l9sl
r
m
iteration 1
m
iteration 2
l,m
iteration 3
r
r
Though binary search is clearly based on a recursive idea, it can he easily
implemented as a nonrecursive algorithm, too. Here is a pseudocode for this
nonrecursive version.
ALGORITHM
BinarySearch(A[O.. n -1], K)
//Implements nonrecursive binary search
//Input: An array A[O .. n - 1] sorted in ascending order and
II
a search key K
//Output: An index of the array's element that is equal to K
II
or -1 if there is no such element
l+-O;r+-n-1
while l <:: r do
m +- L(l + r)/2J
ifK =A[m]returnm
elseifK <A[m] r <--m-1
else! <--m+1
return -1
The standard way to analyze the efficiency of binary search is to count the
number of times the search key is compared with an element of the array. Moreover, for the sake of simplicity, we will count the so-called three-way comparisons.
This assumes that after one comparison of K with A[m], the algorithm can determine whether K is smaller, equal to, or larger than A[m ].
How many such comparisons does the algorithm make on an array of n
elements? The answer obviously depends not only on n but also on the specifics of
a particular instance of the problem. Let us find the number of key comparisons
in the worst case C'""'·"(n). The worst-case inputs include all arrays that do not
contain a given search key (and, in fact, some cases of successful searches as well).
Since after one comparison the algorithm faces the same situation but for an array
half the size, we get the following recurrence relation for Cwa.·"(n):
Cwm·.«(n) = Cw 0 ,."(lnj2J) + 1 for n > 1, Cw"'""(1) = 1.
(4.2)
4.3
Binary Search
137
(Stop and convince yourself that 11j2 must be, indeed, rounded down and that the
initial condition must be written as specified.)
As we discussed in Section 2.4, the standard way of solving recurrences such
as recurrence (4.2) is to assume that 11 = 2k and solve the resulting recurrence
by backward substitutions or another method. We leave this as a straightforward
exercise to obtain the solution
(4.3)
Actually, one can prove that the solution given by formula (4.3) for 11 = 2k can be
tweaked to get a solution valid for an arbitrary positive integer 11:
(4.4)
Let us verify by substitution that cwm·_,(11) = Llogz nJ + 1 indeed satisfies
equation (4.2) for any positive even number 11. (You are asked to do this for odd
11's in Exercises 4.3). If n is positive and even, n = 2i where i > 0. The left-hand
side of equation ( 4.2) for 11 = 2i is
cwm-,(11) = L1og2 11J + 1 = Llog2 2iJ + 1 = Llog2 2 + log2 iJ + 1
= (1 + Llog2 ij) + 1 = Llog2 iJ + 2.
The right-hand side of equation (4.2) for 11 = 2i is
cwm-,1 Cln/2J) + 1 = cwm·,(L2if2J) + 1 = cwm-,1(i) + 1
= (Llog2 ij + 1) + 1 = Llog2 ij + 2.
Since both expressions are the same, we proved the assertion.
Formula ( 4.4) deserves attention. First, it implies that the worst-case efficiency
of binary search is in E>(log n). (Incidentally, we could get this fact by applying the
Master Theorem, but this approach would not give us the value of the multiplicative constant.) Second, it is the answer we should have fully expected: since the
algorithm simply reduces the size of the remaining array by about half on each
iteration, the number of such iterations needed to reduce the initial size 11 to the
final size 1 has to be about log2 11. Third, to reiterate the point made in Section 2.1,
the logarithmic function grows so slowly that its values remain small even for very
large values of 11. In particular, according to formula (4.4), it will take no more than
Llog2 103J + 1 = 10 three-way comparisons to find an element of a given value (or
establish that there is no such element) in any sorted array of 1000 elements, and
it will take no more than Llog2 106 j + 1 = 20 comparisons to do this for any sorted
array of size one million!
What can we say about the average-case efficiency of binary search? A sophisticated analysis shows that the average number of key comparisons made by
binary search is only slightly smaller than that in the worst case:
138
Divide-and-Conquer
(More accurate formulas for the average number of comparisons in a successful
and an unsuccessful search are C~~~ (n) "" log2 n - 1 and C;:~ (n) ""log2 (n + 1),
respectively.)
Though binary search is an optimal searching algorithm if we restrict our operations only to comparisons between keys (see Section 10.2), there are searching
algorithms (see interpolation search in Section 5.6 and hashing in Section 7.3) with
a better average-case efficiency, and one of them (hashing) does not even require
the array to be sorted! These algorithms do require some special calculations in
addition to key comparisons, however. Finally, the idea behind binary search has
several applications beyond searching (see, e.g., [BenOO]). In addition, it can be
applied to solving nonlinear equations in one unknown; we discuss this continuous
analogue of binary search, called the method of bisection, in Section 12.4.
Before we leave this section, one other remark about binary search needs
to be made. Binary search is sometimes presented as the quintessential example
of a divide-and-conquer algorithm. This interpretation is flawed because, in fact,
binary search is a very atypical case of divide-and-conquer. Indeed, according
to the definition given at the beginning of this chapter, the divide-and-conquer
technique divides a problem into several subproblems, each of which needs to
be solved. That is not the case for binary search, where only one of the two
subproblems needs to be solved. Therefore, if binary search is to be considered
as a divide-and-conquer algorithm, it should be looked on as a degenerate case
of this technique. As a matter of fact, binary search fits much better into the class
of decrease-by-half algorithms, which we discuss in Section 5.5. Why then is this
discussion of binary search in this chapter? Partly because of tradition and partly
because a bad example can sometimes make a point that a good example cannot.
-----Exercises 4 . 3 - - - - - - - - - - - - - - 1. a. What is the largest number of key comparisons made by binary search in
searching for a key in the following array?
b. List all the keys of this array that will require the largest number of key
comparisons when searched for by binary search.
c. Find the average number of key comparisons made by binary search in a
successful search in this array. (Assume that each key is searched for with
the same probability.)
d. Find the average number of key comparisons made by binary search in an
unsuccessful search in this array. (Assume that searches for keys in each of
the 14 intervals formed by the array's elements are equally likely.)
~
I
I
4.4
Binary Tree Traversals and Related Properties
139
2. Solve the recurrence Cwm.,(n) = Cwm·.,(ln/2J) + lfor n > 1, Cwm-,(1) = 1, for
n = 2k by backward substitutions.
3. a. Prove the equality
[logz nJ + 1 = rlogz(n + 1)1
for n 2: 1.
b. Prove that Cwm-,(n) = [log2 nj + 1 satisfies equation (4.2) for every positive odd integer n.
4. Estimate how many times faster an average successful search will be in a
sorted array of 100,000 elements if it is done by binary search versus sequential
search.
5. Sequential search can be used with about the same efficiency whether a list is
implemented as an array or as a linked list. Is it also true for binary search?
(Of course, we assume that a list is sorted for binary search.)
6. How can one use binary search for range searching, i.e., for finding all the
elements in a sorted array whose values fall between two given values L and
U (inclusively), L :5: U? What is the worst-case efficiency of this algorithm?
7. Write a pseudocode for a recursive version of binary search.
8. Design a version of binary search that uses only two-way comparisons such
as :5: and =. Implement your algorithm in the language of your choice and
carefully debug it (such programs are notorious for being prone to bugs).
9. Analyze the time efficiency of the two-way comparison version designed in
Problem 8.
10. Picture guessing A version of the popular problem-solving task involves
presenting people with an array of 42 pictures-seven rows of six pictures
each-and asking them to identify the target picture by asking questions that
can be answered yes or no. Further, people are then required to identify the
picture with as few questions as possible. Suggest the most efficient algorithm
for this problem and indicate the largest number of questions that may be
necessary.
4.4 Binary Tree Traversals and
Related Properties
In this section, we see how the divide-and-conquer technique can be applied to
binary trees. A binary tree T is defined as a finite set of nodes that is either empty
or consists of a root and two disjoint binary trees TL and T11 called, respectively, the
left and right subtree of the root. We usually think of a binary tree as a special case
of an ordered tree (Figure 4.4). (This standard interpretation was an alternative
definition of a binary tree in Section 1.4.)
140
Divide-and-Conquer
FIGURE 4.4 Standard representation of a binary tree
Since the definition itself divides a binary tree into two smaller structures of
the same type, the left subtree and the right subtree, many problems about binary
trees can be solved by applying the divide-conquer technique. As an example, let
us consider a recursive algorithm for computing the height of a binary tree. Recall
that the height is defined as the length of the longest path from the root to a leaf.
Hence, it can be computed as the maximum of the heights of the root's left and
right subtrees plus 1. (We add 1 to account for the extra level of the root.) Also
note that it is convenient to define the height of the empty tree as -1. Thus, we
have the following recursive algorithm.
ALGORITHM
Height(T)
//Computes recursively the height of a binary tree
//Input: A binary tree T
//Output: The height ofT
ifT = 0 return -1
else return max(Height(TL), Height(TR)} + 1
We measure the problem's instance size by the number of nodes n(T) in a
given binary tree T. Obviously, the number of comparisons made to compute
the maximum of two numbers and the number of additions A(n(T)) made by the
algorithm are the same. We have the following recurrence relation for A(n(T)):
A(n(T)) = A(n(TL)) + A(n(TR)) + 1
for n(T) > 0,
A(O) = 0.
Before we solve this recurrence (can you tell what its solution is?), let us note
that addition is not the most frequently executed operation of this algorithm. What
is? Checking-and this is very typical for binary tree algorithms-that the tree is
not empty. For example, for the empty tree, the comparison T = 0 is executed
once but there are no additions, and for a single-node tree, the comparison and
addition numbers are three and one, respectively.
r
4.4
Binary Tree Traversals and Related Properties
(a)
141
(b)
FIGURE 4.5 (a) Binary tree. (b) Its extension. Internal nodes are shown as circles;
external nodes are shown as squares.
It helps in the analysis of tree algorithms to draw the tree's extension by
replacing the empty subtrees by special nodes. The extra nodes (shown by little
squares in Figure 4.5) are called external; the original nodes (shown by little
circles) are called internal. By definition, the extension of the empty binary tree
is a single external node.
It is easy to see that the height algorithm makes exactly one addition for
every internal node of the extended tree, and it makes one comparison to check
whether the tree is empty for every internal and external node. Thus, to ascertain
the algorithm's efficiency, we need to know how many external nodes an extended
binary tree with n internal nodes can have. Checking Figure 4.5 and a few similar
examples, it is easy to hypothesize that the number of external nodes x is always
one more than the number of internal nodes n:
x=n+l.
(4.5)
To prove this formula, consider the total number of nodes, both internal and
external. Since every node, except the root, is one of the two children of an internal
node, we have the equation
2n + l=x +n,
which immediately implies equation (4.5).
Note that equation (4.5) also applies to any nonempty full binary tree, in
which, by definition, every node has either zero or two children: for a full binary
tree, n and x denote the numbers of parental nodes and leaves, respectively.
Returning to algorithm Height, the number of comparisons to check whether
the tree is empty is
C(n)=n+x=2n+l,
while the number of additions is
A(n) = n.
i
142
Divide-and-Conquer
a
c
b
e
d
f
Preorder:
Inorder:
Postorder:
a, b, d, g, e, c, f
d, g, b, e, a, f, c
g, d, e, b, f, c, a
g
FIGURE 4.6 Binary tree and its traversals
The most important divide-and-conquer algorithms for binary trees are the
three classic traversals: preorder, inorder, and postorder. All three traversals visit
nodes of a binary tree recursively, i.e., by visiting the tree's root and its left and
right subtrees. They differ just by the timing of the root's visit:
In the preorder traversal, the root is visited before the left and right subtrees
are visited (in that order).
In the inorder traversal, the root is visited after visiting its left subtree but
before visiting the right subtree.
In the postorder traversal, the root is visited after visiting the left and right
subtrees (in that order).
These traversals are illustrated in Figure 4.6. Their pseudocodes are quite
straightforward, repeating the descriptions given above. (These traversals are also
a standard feature of data structures textbooks.) As to their efficiency analysis, it
is identical to the above analysis of the Height algorithm because a recursive call
is made for each node of an extended binary tree.
Finally, we should note that, obviously, not all questions about binary trees
require traversals of both left and right subtrees. For example, the find and insert
operations for a binary search tree require processing only one of the two subtrees.
Hence, they should be considered not as applications of divide-and-conquer but
rather as examples of the variable-size decrease techuique discussed in Section 5.6.
-----Exercises 4 . 4 - - - - - - - - - - - - - - - 1. Design a divide-and-conquer algorithm for computing the number of levels
in a binary tree. (In particular, the algorithm must return 0 and l for the
empty and single-node trees, respectively.) What is the efficiency class of your
algorithm?
4.4
Binary Tree Traversals and Related Properties
143
2. The following algorithm seeks to compute the number of leaves in a binary
tree.
ALGORITHM
LeafCounter(T)
//Computes recursively the number of leaves in a binary tree
//Input: A binary tree T
//Output: The number of leaves in T
if T = 0 return 0
else return LeafCounter(T1) + LeafCounter(TR)
Is this algorithm correct? If it is, prove it; if it is not, make an appropriate
correction.
3. Prove equality (4.5) by mathematical induction.
4. Traverse the following binary tree
a. in preorder.
b. in inorder.
c. in postorder.
a
b
d
c
e
f
5. Write a pseudocode for one of the classic traversal algorithms (preorder,
inorder, and postorder) for binary trees. Assuming that your algorithm is
recursive, find the number of recursive calls made.
6. Which of the three classic traversal algorithms yields a sorted list if applied to
a binary search tree? Prove this property.
7. a. Draw a binary tree with ten nodes labeled 0, 1, 2, ... , 9 in such a way that
the inorder and postorder traversals of the tree yield the following lists:
9, 3, 1, 0, 4, 2, 7, 6, 8, 5 (inorder) and 9, 1, 4, 0, 3, 6, 7, 5, 8, 2 (postorder).
b. Give an exampleoftwopermutationsofthesamen labelsO, 1, 2, ... , n -1
that cannot be in order and postorder traversal lists of the same binary tree.
c. Design an algorithm that constructs a binary tree for which two given lists
of n labels 0, 1, 2, ... , n - 1 are generated by the inorder and postorder
traversals of the tree. Your algorithm should also identify inputs for which
the problem has no solution.
8. The internal path length I of an extended binary tree is defined as the sum
of the lengths of the paths-taken over all internal nodes-from the root to
I
h
I
I
r
''
'
144
Divide-and-Conquer
each internal node. Similarly, the external path length E of an extended binary
tree is defined as the sum of the lengths of the paths-taken over all external
nodes-from the root to each external node. Prove that E = I + 2n where n
is the number of internal nodes in the tree.
9. Write a program for computing the internal path length of an extended binary
tree. Use it to investigate empirically the average number of key comparisons
for searching in a randomly generated binary search tree.
10. Chocolate bar puzzle Given an n-by-m chocolate bar, you need to break it
into nm r-by-r pieces. You can break a bar only in a straight line, and only one
bar can be broken at a time. Design an algorithm that solves the problem with
the minimum number of bar breaks. What is this minimum number? Justify
your answer by using properties of a binary tree.
4.5 Multiplication of large Integers and
Strassen's Matrix Multiplication
In this section, we examine two surprising algorithms for seemingly straightforward tasks: multiplying two numbers and multiplying two square matrices. Both
seek to decrease the total number of multiplications performed at the expense of
a slight increase in the number of additions. Both do this by exploiting the divideand-conquer idea.
Multiplication of large Integers
Some applications, notably modern cryptology, require manipulation of integers
that are over roo decimal digits long. Since such integers are too long to fit in a
single word of a modern computer, they require special treatment. This practical
need supports investigations of algorithms for efficient manipulation of large
integers. In this section, we outline an interesting algorithm for multiplying such
numbers. Obviously, if we use the classic pen-and-pencil algorithm for multiplying
two n-digit integers, each of the n digits of the first number is multiplied by each
of the n digits of the second number for the total of n 2 digit multiplications. (If
one of the numbers has fewer digits than the other, we can pad a shorter number
with leading zeros to equal their lengths.) Though it might appear that it would be
impossible to design an algorithm with fewer than n2 digit multiplications, it turns
out not to be the case. The miracle of divide-and-conquer comes to the rescue to
accomplish this feat.
To demonstrate the basic idea of the algorithm, let us start with a case of
two-digit integers, say, 23 and 14. These numbers can be represented as follows:
23 = 2. ro 1 + 3. ro0 and r4 = r. 101 + 4. 10°.
4.5
Multiplication of Large Integers and Strassen's Matrix Multiplication
145
Now let us multiply them:
23 d4 = (2 . 101 + 3. 10°) * (1 . 101 + 4. 10°)
2
= (2 * 1)10 + (2 * 4 + 3 * 1)101 + (3 * 4)10°.
The last formula yields the correct answer of 322, of course, but it uses the same
four digit multiplications as the pen-and-pencil algorithm. Fortunately, we can
compute the middle term with just one digit multiplication by taking advantage
of the products 2 * 1 and 3 * 4 that need to be computed anyway:
2 * 4 + 3 * 1 = (2 + 3) * (1 + 4) - 2 * 1- 3 * 4.
Of course, there is nothing special about the numbers we just multiplied.
For any pair of two-digit integers a = a1a0 and b = b 1b0 , their product c can be
computed by the formula
c =a* b = c2 102 + c1101 + c0 ,
where
c2 = a1 * b1 is the product of their first digits,
co = ao * bo is the product of their second digits,
c1 = (at+ ao) * (bt + bo) - (c2 +co) is the product of the sum of the a's digits
and the sum of the b's digits minus the sum of c2 and c0 .
Now we apply this trick to multiplying two n-digit integers a and b where n is
a positive even number. Let us divide both numbers in the middle-after all, we
promised to take advantage of the divide-and-conquer technique. We denote the
first half of the a's digits by a 1 and the second half by a0 ; forb, the notations are b1
and b 0 , respectively. In these notations, a= a 1a 0 implies that a= a 110"12 + a0 , and
b = btbo implies that b = b 110"i2 + b0 . Therefore, taking advantage of the same
trick we used for two-digit numbers, we get
c =a* b = (at10"/2 + ao) * (bt10"i2 + bo)
= (at* ht)10" + (at* bo + ao * b1)10"/2 + (a0 * b0 )
= czlO" + C[ 10"12 +Co,
where
cz = a1 * b1 is the product of their first halves,
co = ao * bo is the product of their second halves,
c1 =(at+ ao) * (bt + bo) - (c2 +co) is the product of the sum of the a's halves
and the sum of the b's halves minus the sum of c2 and c0 .
If n/2 is even, we can apply the same method for computing the products c2,
c0 , and c1. Thus, if n is a power of 2, we have a recursive algorithm for computing
the product of two n-digit integers. In its pure form, the recursion is stopped when
146
Divide-and-Conquer
n becomes one. It can also be stopped when we deem n small enough to multiply
the numbers of that size directly.
How many digit multiplications does this algorithm make? Since multiplication of n-digit numbers requires three multiplications of n/2-digit numbers, the
recurrence for the number of multiplications M (n) will be
M(n) = 3M(n/2)
for n > 1,
M(l) = 1.
Solving it by backward substitutions for n = 2k yields
M(2k) = 3M(2k-l) = 3[3M(2k- 2 )] = 32 M(2k- 2 )
= ... = 3' M(2k-i) = ... = 3k M(2k-k) = 3k
Since k = log2 n,
(On the last step, we took advantage of the following property of logarithms:
aiogh c = clogb a.)
You should keep in mind that for moderately large integers, this algorithm will
probably run longer than the classic one. Brassard and Bratley ([Bra96], pp. 7071) report that in their experiments the divide-and-conquer algorithm started to
outperform the pen-and-pencil method on integers over 600 digits long. If you
program in an object-oriented language such as Java, C++, or Smalltalk, you
should also be aware that these languages have special classes for dealing with
large integers.
Strassen's Matrix Multiplication
Now that we have seen that the divide-and-conquer approach can reduce the
number of one-digit multiplications in multiplying two integers, we should not be
surprised that a similar feat can be accomplished for multiplying matrices. Such
an algorithm was published by V. Strassen in 1969 [Str69]. The principal insight
of the algorithm lies in the discovery that we can find the product C of two 2-by-2
matrices A and B with just seven multiplications as opposed to the eight required
by the brute-force algorithm (see Example 3, Section 2.3). This is accomplished
by using the following formulas:
=
-
bOl]
["oo
"m] *[boo
aw
bw bn
a11
[ m1+~-~+m7
+
m2
m4
m3+~
m1
+ m3 - m 2 + m 6
]
'
4.5
Multiplication of Large Integers and Strassen's Matrix Multiplication
147
where
m1 = (aoo +au)
* (boo+ b1 tl
mz = (aw +au)* boo
m3 ="no* (boJ -
m 4 = a 11
h11l
* (b10- h0o)
ms = (aoo + "tnl
* bu
m6 = (aw- "ool
* (boo+ h01)
m7 = (aot -au)* (bto
+ bu).
Thus, to multiply two 2-by-2 matrices, Strassen's algorithm makes seven multiplications and 18 additions/subtractions, whereas the brute-force algorithm requires
eight multiplications and four additions. These numbers should not lead us to multiplying 2-by-2 matrices by Strassen's algorithm. Its importance stems from its
asymptotic superiority as matrix order n goes to infinity.
Let A and B be two n-by-n matrices where n is a power of two. (If n is not a
power of two, matrices can be padded with rows and columns of zeros.) We can
divide A, B, and their product C into four n /2-by-n /2 submatrices each as follows:
[
Coo
C01] = [~] * [--+---Boo
B01]
Cto
C 11
A 10 I A 11
B10
Bu
It is not difficult to verify that one can treat these submatrices as numbers to get the
correct product. For example, C00 can be computed either as A00 * B00 + A01 * Bw
or as M 1 + M4 - Ms + M 7 where M1, M4, M 5, and M7 are found by Strassen's
formulas, with the numbers replaced by the corresponding submatrices. If the
seven products of n /2-by-n /2 matrices are computed recursively by the same
method, we have Strassen's algorithm for matrix multiplication.
Let us evaluate the asymptotic efficiency of this algorithm. If M (n) is the
number of multiplications made by Strassen's algorithm in multiplying two n-by-n
matrices (where n is a power of2), we get the following recurrence relation for it:
M(n)=7M(n/2)
forn>l,
M(l)=l.
M(2k) = 7M(2k-l) = 7[7 M(2k- 2 )] = 72 M(2k- 2 ) = ...
= 7' M(2k-i) ... = 7k M(2k-k) = 7k
Since k = log2 n,
M(n) = 7log2 n = 11 log2 7 ~ 11 2.807,
which is smaller than n 3 required by the brute-force algorithm.
Since this saving in the number of multiplications was achieved at the expense
of making extra additions, we must check the number of additions A (n) made by
ii
n.··
.·•· .rr.
.'i
148
Divide-and-Conquer
Strassen's algorithm. To multiply two matrices of order 11 > 1, the algorithm needs
to multiply seven matrices of order 11/2 and make 18 additions of matrices of size
n/2; when n = 1, no additions are made since two numbers are simply multiplied.
These observations yield the following recurrence relation:
A(11) = 7 A(n/2) + 18(11/2) 2
for 11 > 1,
A(1) = 0.
Though one can obtain a closed-form solution to this recurrence (see Problem
8), here we simply establish the solution's order of growth. According to the
Master Theorem stated in the beginning of the chapter, A(n) E 8(n 10g, 7). In other
words, the number of additions has the same order of growth as the number
of multiplications. This puts Strassen's algorithm in 8(11 10g2 7 ), which is a better
efficiency class than 8(n 3) of the brute-force method.
Since the time of Strassen's discovery, several other algorithms for multiplying
two 11-by-n matrices of real numbers in O(n") time with progressively smaller
constants a have been invented. The fastest algorithm so far is that of Coopersmith
and Winograd [Coo87], with its efficiency in 0(11 2376 ). The decreasing values of
the exponents have been obtained at the expense of increasing complexity of these
algorithms. Because of large multiplicative constants, none of them is of practical
value. However, they are interesting from a theoretical point of view. Although
these algorithms get closer and closer to the best theoretical lower bound known
for matrix multiplication, which is n 2 multiplications, the gap between this bound
and the best available algorithm remains unresolved. It is also worth mentioning
that matrix multiplication is known to be computationally equivalent to some
other important problems such as solving systems of linear equations.
-----Exercises 4 . 5 - - - - - - - - - - - - - - - 1. Wbat are the smallest and largest numbers of digits the product of two decimal
n-digit integers can have?
2. Compute 2101 * 1130 by applying the divide-and-conquer algorithm outlined
in the text.
3. a. Prove the equality a log, c = c10gb a, which was used twice in Section 4.5.
b. Wby is 11logz 3 better than 3Iog,, as a closed-form formula for M(n)?
4. a. Why did we not include multiplications by 10" in the multiplication count
M(11) of the large-integer multiplication algorithm?
b. In addition to assuming that 11 is a power of 2, we made, for the sake of
simplicity, another, more subtle, assumption in setting up a recurrence
relation for M (11), which is not always true (it does not change the final
answer, however). What is this assumption?
4.6
Closest-Pair and Convex-Hull Problems by Divide-and-Conquer
149
5. How many one-digit additions are made by the pen-and-pencil algorithm in
multiplying two n-digit integers? (You may disregard potential carries.)
6. Verify the formulas underlying Strassen's algorithm for multiplying 2-by-2
matrices.
7. Apply Strassen's algorithm to compute
1 0 2 1] [0 1
1 0
0 41]
[
4 1 1 0
2
0130*2 0 1 1
5 0 2 1
1 3 5 0
exiting the recursion when n = 2, i.e., computing the products of 2-by-2 matrices by the brute-force algorithm.
8. Solve the recurrence for the number of additions required by Strassen's algorithm. (Assume that n is a power of 2.)
9. V. Pan [Pan78] has discovered a divide-and-conquer matrix multiplication
algorithm that is based on multiplying two 70-by-70 matrices using 143,640
multiplications. Find the asymptotic efficiency of Pan's algorithm (you may
ignore additions) and compare it with that of Strassen's algorithm.
10. Practical implementations of Strassen 's algorithm usually switch to the bruteforce method after matrix sizes become smaller than some "crossover point."
Run an experiment to determine such crossover point on your computer
system.
4.6 Closest-Pair and Convex-Hull Problems
by Divide-and-Conquer
In Section 3.3, we discussed the brute-force approach to solving two classic problems of computational geometry: the closest-pair problem and the convex-hull
problem. We saw that the two-dimensional versions of these problems can be
solved by brute-force algorithms in 8 (n 2 ) and 0 (n 3 ) time, respectively. In this section, we discuss more sophisticated and asymptotically more efficient algorithms
for these problems, which are based on the divide-and-conquer technique.
Closest-Pair Problem
Let Pt = (xt, y1), ... , Pn = (xn, y,) be a setS of n points in the plane, where n,
for simplicity, is a power of two. With no loss of generality, we can assume that
the points are ordered in ascending order of their x coordinates. (If they were not,
we can sort them in O(n log n) time, e.g., by mergesort.) We can divide the points
150
Divide-and-Conquer
given into two subsets S1 and S2 of 11(2 points each by drawing a vertical line x = c
so that n/2 points lie to the left of or on the line itself, and 11/2 points lie to the
right of or on the line. (One way of finding an appropriate value for constant c for
doing this is to use the median I"' of the x coordinates.)
Following the divide-and-conquer approach, we can find recursively the closest pairs for the left subset S1 and the right subset S2. Let d1 and d2 be the smallest
distances between pairs of points in S1 and S2, respectively, and let d = min{d1o
d2 ). Unfortunately, d is not necessarily the smallest distance between all pairs of
points in S1 and S2 because a closer pair of points can lie on the opposite sides
of the separating line. So, as a step of combining the solutions to the smaller subproblems, we need to examine such points. Obviously, we can limit our attention
to the points in the symmetric vertical strip of width 2d since the distance between
any other pair of points is greater than d (Figure 4.7a).
Let C1 and C2 be the subsets of points in the left and right parts of the strip,
respectively. Now, for every point P(x, y) in C1o we need to inspect points in C2 that
may be closer toP than d. Obviously, such points must have their y coordinates in
the interval [y - d, y + d]. The critical insight here is an observation that there can
be no more than six such points because any pair of points in C2 is at least d apart
from each other. (Recall that d ::: d2 where d2 is the smallest distance between
pairs of points to the right of the dividing line.) The worst case is illustrated in
Figure 4.7b.
Another important observation is that we can maintain lists of points in C1
and C2 sorted in ascending order of their y coordinates. (You can think of these
lists as projections of the points on the dividing line.) Moreover, this ordering can
be maintained not by resorting points on each iteration but rather by merging two
previously sorted lists (see algorithm Merge in Section 4.1). We can process the C1
points sequentially while a pointer into the C2 list scans an interval of width 2d to
fetch up to six candidates for computing their distances to a current point P of the
C1 list. The time M(11) for this "merging" of solutions to the smaller subproblems
is in 0(11).
We have the following recurrence for T (n), the running time of this algorithm
on n presorted points:
T(n) = 2T(11/2) + M(n).
Applying the 0 version of the Master Theorem (with a= 2, b = 2, and d = 1),
we get T (n) E 0 (11 log 11). The possible necessity to presort input points does not
change the overall efficiency class if sorting is done by a O(n log n) algorithm. In
fact, this is the best efficiency class we can achieve because it has been proved that
any algorithm for this problem must be in i:2(11log 11) (see [Pre85], p. 188).
Convex-Hull Problem
Let us revisit the convex-hull problem introduced in Section 3.3: fmd the smallest
convex polygon that contains 11 given points in the plane. We consider here a
4.6 Closest-Pair and Convex-Hull Problems by Divide-and-Conquer
X= C
•
•
•
•
151
"k
•
•
•
•
d
d
•
(a)
d
d
P,
d
(b)
FIGURE 4.7 (a) Idea of the divide-and-conquer algorithm for the closest-pair problem.
(b) The six points that may need to be examined for point P.
divide-and-conquer algorithm called quickhull because of its resemblance to
quicksort.
Let P1 = (xJ, y 1), ... , P, = (x,, y,) be a setS of n > 1 points in the plane. We
assume that the points are sorted in increasing order of their x coordinates, with
ties resolved by increasing order of the y coordinates of the points involved. It
is not difficult to prove the geometrically obvious fact that the leftmost point P1
and the rightmost point P, are two distinct extreme points of the set's convex hull
Divide-and-Conquer
•
•
•
•
•
•
•
FIGURE 4.8 Upper and lower hulls of a set of points
(Figure 4.8). Let P;P n be the straight line through points P1 and Pn directed from
P1 to P.. Tbis line separates the points of S into two sets: S1 is the set of points to
the left of or on this line and S2 is the set of points to the right of or on this line.
(We say that point p 3 is to the left of the line P1P2 directed from point p 1 to point
P2 if PtPzP 3 forms a counterclockwise cycle. Later, we cite an analytical way to
check this condition based on checking the sign of a determinant formed by the
coordinates of the three points.) The points of S on the line P;P "' other than P1
and Pn, cannot be extreme points of the convex hull and hence are excluded from
further consideration.
The boundary of the convex hull of S is made up of two polygonal chains:
an "upper" boundary and a "lower" boundary. The "upper" boundary, called the
upper hull, is a sequence of line segments with vertices at P1, some of the points
in S1 (if S1 is not empty), and Pn- The "lower" boundary, called the lower hull, is
a sequence of line segments with vertices at P1, some of the points in S2 (if S2 is
not empty) and PnThe fact that the convex hull of the entire set S is composed of the upper and
lower hulls, which can be constructed independently and in a similar fashion, is a
very useful observation that is exploited by several algorithms for this problem.
For concreteness, let us discuss how quickhull proceeds to construct the upper
hull; the lower hull can be constructed in the same manner. If S1 is empty, the upper
hull is simply the line segment with the endpoints at P 1 and P.. If S1 is not empty,
the algorithm identifies vertex p max in sj, which is the farthest from the line P;P n
(Figure 4.9) If there is a tie, the point that maximizes the angle L PmaxPJ Pn can
be selected. (Note that point Pmax maximizes the area of the triangle with two
vertices at P1 and Pn and the third at some other point of SJ-) Then the algorithm
identifies all the points of set S1 that are to the left of the line P;P max; these are
the points that, along with P1 and Pmax• will make up the set St.!· The points of St
4.6
Closest-Pair and Convex-Hull Problems by Divide-and-Conquer
•
•
•
153
•
•
•
FIGURE 4.9 The idea of quickhull
to the left ofthe line p;;:;;P, will make up, along with Pmax and P,. the set S1.2. It
is not difficult to prove that
ill
ill
.,
Pmax is a vertex of the upper hull;
the points inside /',. P1 Pmax P, cannot be vertices of the upper hull (and hence
can be eliminated from further consideration); and
there are no points to the left of both lines P;P max and p;;:;;P,.
Therefore, the algorithm can continue constructing the upper hulls of ? 1 U Su U
Pmax and Pmax U S1.2 U P, recursively and then simply concatenate them to getthe
upper hull of the entire set P 1 U S1 UP,.
Now we have to figure out how the algorithm's geometric operations can be
implemented. Fortunately, we can take advantage of the following very useful fact
from analytical geometry: if p 1 = (x1, y 1), p2 = (x 1, y 1), and p 3 = (x3 , y 3 ) are three
arbitrary points in the Cartesian plane, then the area of the triangle !',.p 1p 2 p 3 is
equal to one half of the magnitude of the determinant
1
1 = XtYz + x3Y1 + XzY3- x3Y2- XzYl- XJY3·
1
while the sign of this expression is positive if and only if the point P3 = (x3, Y3)
is to the left of the line p;P,. Using this formula, we can cheek in constant time
whether a point lies to the left of the line determined by two other points as well
as find the distance from the point to the line.
Quickhull has the same 8 (n 2 ) worst -case efficiency as quicksort (Problem 8
in the exercises). In the average case, however, we should expect a much better
performance. First, the algorithm should benefit from the quicksort-like savings
from the on-average balanced split of the problem into two smaller subproblems.
Second, a siginificant fraction of the points-namely those inside b. P1 PmaxPn (see
154
Divide-and-Conquer
Figure 4.9)-are eliminated from further processing. Under a natural assumption
that points given are chosen randomly from a uniform distribution over some
convex region (e.g., a circle or a rectangle), the average-case efficiency of quickhull
turns out to be linear [Ove80].
-----Exercises 4J5 - - - - - - - - - - - - - - - 1. a. For the one-dimensional version of the closest-pair problem, i.e., for the
problem of finding two closest numbers among a given set of n real numbers, design an algorithm that is directly based on the divide-and-conquer
technique and determine its efficiency class.
b. Is it a good algorithm for this problem?
2. Consider the version of the divide-and-conquer two-dimensional closest-pair
algorithm in which we simply sort each of the two sets C1 and C2 in ascending
order of their y coordinates on each recursive call. Assuming that sorting is
done by mergesort, set up a recurrence relation for the running time in the
worst case and solve it for n = 2k
3. Implement the divide-and-conquer closest-pair algorithm, outlined in this
section, in the language of your choice.
4. Find on the Web a visualization of an algorithm for the closest-pair problem.
What algorithm does this visualization represent?
5. The Voronoi polygon for a point P of a setS of points in the plane is defined
to be the perimeter of the set of all points in the plane closer to P than to any
other point in S. The union of all the Voronoi polygons of the points in S is
called the Voronoi diagram of S.
a. What is the Voronoi diagram for a set of three points?
b. Find on the Web a visualization of an algorithm for generating the Voronoi
diagram and study a few examples of such diagrams. Based on your observations, can you tell how the solution to the previous question is generalized to the general case?
6. Explain how one can find point Pmax in the quickhull algorithm analytically.
i7,, What is the best-case efficiency of quickhull?
\
·s:; Give a specific example of inputs that make the quickhull algorithm run in
quadratic time.
9. Implement the quickhull algorithm in the language of your choice.
10. Shortest path around There is a fenced area in the two-dimensional Euclidean plane in the shape of a convex polygon with vertices at points P1 (x 1 , y 1),
P2(x 2, y 2), ... , PnCxn, Ynl (not necessarily in this order). There are two more
Summary
155
points, A(xA, YA) and B(x 8 , y 8 ), such that xA < min{x1 , x 2 , ... , x,) and x 8 >
max{x 1, x 2 , ... , x,). Design a reasonably efficient algorithm for computing
the length of the shortest path between A and B. [0Ro98], p. 68
SUMMARY
"
Divide-and-conquer is a general algorithm design technique that solves a
problem's instance by dividing it into several smaller instances (ideally,
of equal size), solving each of them recursively, and then combining their
solutions to get a solution to the original instance of the problem. Many
efficient algorithms are based on this technique, although it can be both
inapplicable and inferior to simpler algorithmic solutions.
"
Running time T(11) of many divide-and-conquer algorithms satisfies the
recurrence T (11) = aT(11 /h) + f (n ). The Master Theorem establishes the order
of growth of its solutions.
'"
Mergesort is a divide-and-conquer sorting algorithm. It works by dividing an
input array into two halves, sorting them recursively, and then merging the two
sorted halves to get the original array sorted. The algorithm's time efficiency
is in 8 (11 log n) in all cases, with the number of key comparisons being very
close to the theoretical minimum. Its principal drawback is a significant extra
storage requirement.
"
Quicksort is a divide-and-conquer sorting algorithm that works by partitioning its input's elements according to their value relative to some preselected
element. Quicksort is noted for its superior efficiency among 11 log n algorithms for sorting randomly ordered arrays but also for the quadratic
worst-case efficiency.
"
Binary search is a 0 (log 11) algorithm for searching in sorted arrays. It is
an atypical example of an application of the divide-and-conquer technique
because it needs to solve just one problem of half the size on each of its
iterations.
11
The classic traversals of a binary tree-preorder, inorder, and postorder-
and similar algorithms that require recursive processing of both left and right
subtrees can be considered examples of the divide-and-conquer technique.
Their analysis is helped by replacing all the empty subtrees of a given tree
with special external nodes.
"
There is a divide-and-conquer algorithm for multiplying two n-digit integers
that requires about 11 t.sss one-digit multiplications.
156
Divide-and-Conquer
,.
Strassen 's algorithm needs only seven multiplications to multiply two 2-by2 matrices but requires more additions than the definition-based algorithm.
By exploiting the divide-and-conquer technique, this algorithm can multiply
two n-by-n matrices with about n2 ·807 multiplications.
"
The divide-and-conquer technique can be successfully applied to two important problems of computational geometry: the closest-pair problem and the
convex-hnll problem.
Decrease-and-Conquer
Plutarch says that Sertorius, in order to teach his soldiers that perseverance
and wit are better than brute force, had two horses brought before them,
and set two men to pull out their tails. One of the men was a burly Hercules,
who tugged and tugged, but all to no purpose; the other was a sharp, weaselfaced tailor, who plucked one hair at a time, amidst roars of laughter, and
soon left the tail quite bare.
-E. Cobham Brewer, Dictionary of Phrase and Fable, 1898
T
he decrease-and-conquer technique is based on exploiting the relationship
between a solution to a given instance of a problem and a solution to a smaller
instance of the same problem. Once such a relationship is established, it can be
exploited either top down (recursively) or bottom up (without a recursion). There
are three major variations of decrease-and-conquer:
"
"
!!Ill
decrease by a constant
decrease by a constant factor
variable size decrease
In the decrease-by-a-constant variation, the size of an instance is reduced by
the same constant on each iteration of the algorithm. Typically, this constant is
equal to one (Figure 5.1), although reduction-by-two cases do happen occasionally, for example, in algorithms that have to act differently for instances of odd
and even sizes.
Consider, as an example, the exponentiation problem of computing a" for
positive integer exponents. The relationship between a solution to an instance of
sizen and an instance of size n- 1 is obtained by the obvious formula: an= an-t. a.
So the function .f (n) =a" can be computed either "top down" by using its recursive
definition
.f(n) = { .f(n- 1) ·a
a
if n > 1
Ifn=1
(5.1)
157
158
Decrease-and-Conquer
problem of size n
subproblem
ofsizen-1
solution to
the subproblem
solution to
the original problem
FIGURE 5.1 Decrease (by one)-and-conquer technique
or "bottom up" by multiplying a by itself n - 1 times. (Yes, it is the same as the
brute-force algorithm, but we have come to it by a different thought process.)
More interesting examples of decrease-by-one algorithms appear in Sections 5.1-
5.4.
The decrease-by-a-constant-factor technique suggests reducing a problem's
instance by the same constant factor on each iteration of the algorithm. In most
applications, this constant factor is equal to two. (Can you give an example of such
an algorithm?) The decrease-by-half idea is illustrated in Figure 5.2.
For an example, let us revisit the exponentiation problem. If the instance of
size n is to compute an, the instance of half its size will be to compute an/2 , with
the obvious relationship between the two: a"= (a"/2 ) 2 But since we consider here
instances of the exponentiation problem with integer exponents only, the former
does not work for odd n. If n is odd, we have to compute a"- 1 byusingtherule for
even-valued exponents and then multiply the result by a. To summarize, we have
the following formula:
!
J
Decrease-and-Conquer
problem
159
of size n
subproblem
of size n/2
solution to
the subproblem
solution to
the original problem
FIGURE 5.2 Decrease (by haiti-and-conquer technique
2 2
(a"1 )
a"=
{
(aln-l)/2) 2 . a
a
if n is even and positive
if n is odd and greater than 1
if n = 1.
(5.2)
If we compute a" recursively according to formula (5.2) and measure the algorithm's efficiency by the number of multiplications, we should expect the algorithm
to be in 0 (log n) because, on each iteration, the size is reduced by at least one half
at the expense of no more than two multiplications.
Note a difference between this algorithm and the one based on the divideand-conquer idea of solving two instances of the exponentiation problem of size
nf2:
ifn>l
ifn=l.
(5.3)
The algorithm based on formula (5.3) is inefficient (why?), whereas the one based
on (5.2) is much faster.
160
Decrease-and-Conquer
A few other examples of decrease-by-a-constant-factor algorithms are given
in Section 5.5 and its exercises. Such algorithms are so efficient, however, that
there are few examples of this kind.
Finally, in the val'iable-size-decrease variety of decrease-and-conquer, a size
reduction pattern varies from one iteration of an algorithm to another. Euclid's
algorithm for computing the greatest common divisor provides a good example
of such a situation. Recall that this algorithm is based on the formula
gcd(m, n) = gcd(11, m mod 11).
Though the arguments on the right-hand side are always smaller than those on the
left-hand side (at least starting with the second iteration of the algorithm), they
are smaller neither by a constant nor by a constant factor. A few other examples
of such algorithms appear in Section 5.6.
5.1
Insertion Sort
In this section, we consider an application of the decrease-by-one technique to
sorting an array A[0 .. 11- 1]. Following the technique's idea, we assume that tbe
smaller problem of sorting the array A[O .. n- 2] has already been solved to give
us a sorted array of size 11 - 1: A[O] :<' ... :<'A [11 - 2]. How can we take advantage
of this solution to the smaller problem to get a solution to the original problem
by taking into account the element A[n -1]? Obviously, all we need is to find an
appropriate position for A[n - 1] among the sorted elements and insert it there.
There are three reasonable alternatives for doing this. First, we can scan the
sorted subarray from left to right until the first element greater than or equal
to A[11 -1] is encountered and then insert A[11 -1] right before that element.
Second, we can scan the sorted subarray from right to left until the first element
smaller than or equal to A[n -1] is encountered and then insert A[n -1] right
after that element. These two alternatives are essentially equivalent; usually, it is
the second one that is implemented in practice because it is better for sorted and
almost-sorted arrays (whf\). The resulting algorithm is called straight insertion
sort or simply insertion sort. The third alternative is to use binary search to find an
appropriate position for A [n - 1] in the sorted portion of the array. The resulting
algorithm is called binary insertion sort. We ask you to implement this idea and
investigate the efficiency of binary insertion sort in the exercises to this section.
Though insertion sort is clearly based on a recursive idea, it is more efficient
to implement this algorithm bottom up, i.e., iteratively. As shown in Figure 5.3,
starting with A[1] and ending with A[n - 1], A[i] is inserted in its appropriate place
among the first i elements of the array that have been already sorted (but, unlike
selection sort, are generally not in their final positions).
Here is a pseudocode of this algorithm.
5.1
161
Insertion Sort
A[O] ::0 ... ::0 A[j] < A[j + 1] ::0 ... ::0 A[i -1] I A[i] ... A[n -1]
smaller than or equal to A[i]
greater than A[i]
FIGURE 5.3 Iteration of insertion sort: A[i] is inserted in its proper position among the
preceding elements previously sorted.
ALGORITHM InsertionSort(A[O .. n - 1])
//Sorts a given array by insertion sort
//lnput: An array A[O .. n- 1] of n orderable elements
//Output: Array A[O .. n - 1] sorted in nondecreasing order
for i <--- 1 to n - 1 do
v <--- A[i]
(\
j<---i-1
while j :>: 0 and A[j] > v do
A[j + 1] <--- A[j]
'2
j<---j-1
A[j+1]<-v
The operation of the algorithm is illustrated in Figure 5.4.
The basic operation of the algorithm is the key comparison A[j] > v. (Why
)not j :>: 0? Because it will almost certainly be faster than the former in an actual
~'<~r 1 computer implementation. Moreover, it is not germane to the algorithm: a better
implementation with a sentinel-see Problem 5 in the exercises-eliminates it
,.-,--~----------altogether.)
The number of key comparisons in this algorithm obviously depends on the
nature of the input. In the worst case, A[j] > v is executed the largest number
of times, i.e., for every j = i - 1, ... , 0. Since v = A[i], it happens if and only if
c\"""
89 I 45 68 90 29 34 17
45 89 I 68 90 29 34 17
45 68 89 I 90 29 34 17
45 68 89 90 I 29 34 17
29 45 68 89 90 I 34 17
29 34 45 68 89 90 I 17
17 29 34 45 68 89 90
FIGURE 5.4 Example of sorting with insertion sort. A vertical bar separates the sorted
part of the array from the remaining elements; the element being inserted
is in bold.
162
Decrease-and-Conquer
A[j] > A[i] for j = i- 1, ... , 0. (Note that we are using the fact that on the ith
iteration of insertion sort all the elements preceding A[i] are the first i elements in
the input, albeit in the sorted order.) Thus, for the worst-case input, we get A[O] >
A[1] (fori= 1), A[1] > A[2] (fori= 2), ... , A[n- 2] > A[n -1] (fori= n -1).
In other words, the worst-case input is an array of strictly decreasing values. The
number of key comparisons for such an input is
n-1 i-l
n-1
'\''\'
'\'·
Cw 0 ,-,(n) = L... L...1 = L... l =
i=l j=O
i=l
(
)
n- 1 n
2
E E>(n ).
2
Thus, in the wo~st case, insertion sort makes exactly the same number of compar-
isons as selection sort (see Section 3.1).
In the best case, the comparison A[j] > v is executed only once on every
iteration of the outer loop. It happens if and only if A[i - 1]::: A[i] for every i =
1, ... , n - 1, i.e., if the input array is already sorted in ascending order. (Though
it "makes sense" that the best case of an algorithm happens when the problem
is already solved, it is not always the case: recall our discussion of quicksort in
Chapter 4.) Thus, for sorted arrays, the number of key comparisons is
n-1
ch,,,,(n) =
L 1 = n- 1 E E>(n).
i=l
t
This very good performance in the best case of sorted arrays is not very useful
by itself, because we cannot expect such convenient inputs. However, almostsorted files arise in a variety of applications, and insertion sort preserves its
excellent performance on such inputs. For example, while sorting an array by
quicksort, we can stop the algorithm's iterations after subarrays become smaller
than some predefined size (say, 10 elements). By that time, the entire array is
almost sorted and we can finish the job by applying insertion sort to it. This
modification typically decreases the total running time of quicksort by about 10%.
c\.cc== ... A rigorous analysis of the algorithm's average-case efficiency is based on
investigating the number of element pairs that are out of order (see Problem 8).
It shows that on randomly ordered arrays, insertion sort makes on average half as
many comparisons as on decreasing arrays, i.e.,
This twice-as-fast average-case performance coupled with an excellent efficiency
on almost-sorted arrays makes insertion sort stand out among its principal com-
petitors among elementary sorting algorithms, selection sort and bubble sort. In
addition, its extension named shellsort, after its inventor D. L. Shell [She59], gives
us an even better algorithm for sorting moderately large files (see Problem 10).
5.1
Insertion Sort
163
-----Exercises 5.1 - - - - - - - - - - - - - - - 1. Ferrying soldiers A detachment of n soldiers must cross a wide and deep
river with no bridge in sight. They notice two 12-year-old boys playing in a
rowboat by the shore. The boat is so tiny, however, that it can only hold two
boys or one soldier. How can the soldiers get across the river and leave the
boys in joint possession of the boat? How many times need the boat pass from
shore to shore?
2. Alternating glasses There are 2n glasses standing next to each other in a row,
the first n of them filled with a soda drink, while the remaining n glasses are
empty. Make the glasses alternate in a filled-empty-filled-empty pattern in the
minimum number of glass moves. (Gar78], p. 7
0000
O'i/0 O'i/0
3. Design a decrease-by-one algorithm for generating the power set of a set of n
elements. (The power set of a set S is the set of all the subsets of S, including
the empty set and S itself.)
4. Apply insertion sort to sort the list E, X, A, M, P, L, E in alphabetical order.
5. a. What sentinel should be put before the first element of an array being
sorted to avoid checking the in-bound condition j 2: 0 on each iteration
of the inner loop of insertion sort?
b. Will the version with the sentinel be in the same efficiency class as the
original version?
6. Is it possible to implement insertion sort for sorting linked lists? Will it have
the same O(n 2 ) efficiency as the array version?
7. Consider the following version of insertion sort.
ALGORITHM
InsertSort2(A[O .. n - 1])
fori <--lton-1do
j<--i-1
while j 2:0 and A(j] > A[j + 1] do
swap( A(}], A(j + 1])
j<--j-1
What is its time efficiency? How is it compared to that of the version given in
the text?
8. Let A[O .. n - 1] be an array of n sortable elements. (For simplicity, you can
assume that all the elements are distinct.) A pair (A(i], A[j]) is called an
inversion if i < j and A[i] > A[j].
164
Decrease-and-Conquer
a. What arrays of size n have the largest number of inversions and what is this
number? Answer the same questions for the smallest number of inversions.
b. Show that the average-case number of key comparisons in insertion sort is
given by the formula
9. Binary insertion sort uses binary search to find an appropriate position to
insert A[i] among the previously sorted A[O] :S ... :s A[i - 1]. Determine the
worst-case efficiency class of this algorithm.
10. Shellsort (more accurately Shell's sort) is an important sorting algorithm that
works by applying insertion sort to each of several interleaving sublists of a
given list. On each pass through the list, the sublists in question are formed
by stepping through the list with an increment h, taken from some predefined
decreasing sequence of step sizes, h1 > ... >hi > ... > 1, which must end with
1. (The algorithm works for any such sequence, though some sequences are
known to yield a better efficiency than others. For example, the sequence 1,
4, 13, 40, 121, ... , used, of course, in reverse, is known to be among the best
for this purpose.)
a. Apply shellsort to the list
S, H, E, L, L, S, 0, R, T, I, S, U, S, E, F, U, L
b. Is shellsort a stable sorting algorithm?
c. Implement shellsort, straight insertion sort, binary insertion sort, mergesort, and quicksort in the language of your choice and compare their performance on random arrays of sizes 102 , 103, 104 , and 105 as well as on
increasing and decreasing arrays of these sizes.
!5.2 Depth-First Search and Breadth-First Search
In the next two sections of this chapter, we deal with very important graph algorithms that can be viewed as applications of the decrease-by-one technique.
We assume familiarity with the notion of a graph, its main varieties (undirected,
directed, and weighted graphs), the two principal representations of a graph (adjacency matrix and adjacency lists), and such notions as graph connectivity and
acyclicity. If needed, a brief review of this material can be found in Section 1.4.
As pointed out in Section 1.3, graphs are interesting structures with a wide
variety of applications. Many graph algorithms require processing vertices or
edges of a graph in a systematic fashion. There are two principal algorithms for
doing such traversals: depth-first search (DFS) and breadth-first search (BFS). In
5.2
Depth-First Search and Breadth-First Search
165
addition to doing their main job of visiting vertices and traversing edges of a graph.
these algorithms have proved to be very useful in investigating several important
properties of a graph.
Depth-First Search
Depth-first search starts visiting vertices of a graph at an arbitrary vertex by marking it as having been visited. On each iteration, the algorithm proceeds to an
unvisited vertex that is adjacent to the one it is currently in. (If there are several such vertices. a tie can be resolved arbitrarily. As a practical matter, which
of the adjacent unvisited candidates is chosen is dictated by the data structure
representing the graph. In our examples, we will always break ties by the alphabetical order of the vertices.) This process continues until a dead end-a vertex
with no adjacent unvisited vertices-is encountered. At a dead end, the algorithm
backs up one edge to the vertex it came from and tries to continue visiting unvisited vertices from there. The algorithm eventually halts after backing up to
the starting vertex, with the latter being a dead end. By then, all the vertices in
the same connected component as the starting vertex have been visited. If unvisited vertices still remain, the depth-first search must be restarted at any one of
them.
It is convenient to use a stack to trace the operation of depth-first search. We
push a vertex onto the stack when the vertex is reached for the first time (i.e., the
visit of the vertex starts), and we pop a vertex off the stack when it becomes a
dead end (i.e., the visit of the vertex ends).
It is also very useful to accompany a depth-first search traversal by constructing the so-called depth-first search forest. The traversal's starting vertex serves
as the root of the first tree in such a forest. Whenever a new unvisited vertex is
reached for the first time, it is attached as a child to the vertex from which it is being
reached. Such an edge is called a tree edge because the set of all such edges forms
a forest. The algorithm may also encounter an edge leading to a previously visited
vertex other than its immediate predecessor (i.e., its parent in the tree). Such an
edge is called a back edge because it connects a vertex to its ancestor, other than
the parent, in the depth-first search forest. Figure 5.5 provides an example of a
depth-first search traversal, with the traversal's stack and corresponding depthfirst search forest shown as well.
Here is a pseudocode of the depth-first search.
ALGORITHM
DFS(G)
//Implements a depth-first search traversal of a given graph
//Input: Graph G = (V, E)
//Output: Graph G with its vertices marked with consecutive integers
//in the order they've been first encountered by the DFS traversal
{.,
;tri
rli
I
166
Decrease-and-Conquer
g)-----------1 h
ee. 2
b5, 3
d3 , 1
f4,4
c2. 5
8 1,6
(a)
ho.7
ig,a
ha. 9
97.1o
(b)
(c)
FIGURE 5.5 Example of a DFS traversal. (a) Graph. (b) Traversal's stack (the first subscript
number indicates the order in which a vertex was visited, i.e., pushed onto
the stack; the second one indicates the order in which it became a deadend, i.e., popped off the stack). (c) DFS forest (with the tree edges shown
with solid lines and the back edges shown with dashed lines).
mark each vertex in V with 0 as a mark of being "unvisited"
count+--- 0
for each vertex v in V do
if v is marked with 0
dfs(v)
dfs(v)
//visits recursively all the unvisited vertices connected to vertex v hy a path
//and numbers them in the order they are encountered
//via global variable count
count +---count + 1; mark v with count
for each vertex w in V adjacent to v do
if w is marked with 0
df<(w)
The brevity of the DFS pseudocode and the ease with which it can be performed by hand may create a wrong impression about the level of sophistication
of this algorithm. To appreciate its true power and depth, you should trace the
algorithm's action by looking not at a graph's diagram but at its adjacency matrix
or adjacency lists. (Try it for the graph in Figure 5.5 or a smaller example.)
How efficient is depth-first search? It is not difficult to see that this algorithm
is, in fact, quite efficient since it takes just the time proportional to the size of the
data structure used for representing the graph in question. Thus, for the adjacency
matrix representation, the traversal's time is in 8(1VI 2 ), and for the adjacency
5.2
Depth-First Search and Breadth-First Search
167
list representation, it is in El(IVI +lEI) where lVI and lEI are the number of the
graph's vertices and edges, respectively.
A DFS forest, which is obtained as a by-product of a DFS traversal, deserves a
few comments, too. To begin with, it is not actually a forest. Rather, we can look at
it as the given graph with its edges classified by the DFS traversal into two disjoint
classes: tree edges and back edges. (No other types are possible for a DFS forest
of an undirected graph.) Again, tree edges are edges used by the DFS traversal to
reach previously unvisited vertices. If we consider only the edges in this class, we
will indeed get a forest. Back edges connect vertices to previously visited vertices
other than their immediate predecessors in the traversal. They connect vertices to
their ancestors in the forest other than their parents.
A DFS traversal itself and the forest-like representation of a graph it provides
have proved to be extremely helpful for the development of efficient algorithms
for checking many important properties of graphs1 Note that the DFS yields two
orderings of vertices: the order in which the vertices are reached for the first
time (pushed onto the stack) and the order in which the vertices become dead
ends (popped off the stack). These orders are qualitatively different, and various
applications can take advantage of either of them.
Important elementary applications of DFS include checking connectivity and
checking acyclicity of a graph. Since DFS halts after visiting all the vertices connected by a path to the starting vertex, checking a graph's connectivity can be done
as follows. Start a DFS traversal at an arbitrary vertex and check, after the algo·
rithm halts, whether all the graph's vertices will have been visited. If they have,
the graph is connected; otherwise, it is not connected. More generally, we can use
DFS for identifying connected components of a graph (how?).
As for checking for a cycle presence in a graph, we can take advantage of the
graph's representation in the form of a DFS forest. If the latter does not have back
edges, the graph is clearly acyclic. If there is a back edge from some vertex u to its
ancestor v (e.g., the back edge from d to a in Figure 5.5c), the graph has a cycle
that comprises the path from v to u via a sequence of tree edges in the DFS forest
followed by the back edge from u to v.
You will find a few other applications of DFS later in the book, although more
sophisticated applications, such as finding articulation points of a graph, are not
included. (A vertex of a connected graph is said to be its articulation point if its
removal with all edges incident to it breaks the graph into disjoint pieces.)
Breadth-First Search
If depth-first search is a traversal for the brave (the algorithm goes as far from
"home" as it can), breadth-first search is a traversal for the cautious. It proceeds in
1.
The discovery of several such applications was an important breakthrough achieved by the two
American computer scientists John Hopcroft and Robert Tarjan in the 1970s. For this and other
contributions, they subsequently won the Turing Award-the most important prize given in theoretical
computer science [Hop87, Tar87].
168
Decrease-and-Conquer
h
g
j
~
4
(a)
a1 c2 d3 e4 fs b5
97 hsjg iw
(b)
j
(c)
FIGURE 5.6 Example of a BFS traversal. (a) Graph. (b) Traversal's queue, with the
numbers indicating the order in which the vertices were visited, i.e., added
to (or removed from) the queue. (c) BFS forest (with the tree edges shown
with solid lines and the cross edges shown with dotted lines).
a concentric manner by visiting first all the vertices that are adjacent to a starting
vertex, then all unvisited vertices two edges apart from it, and so on, until all
the vertices in the same connected component as the starting vertex are visited.
If there still remain unvisited vertices, the algorithm has to be restarted at an
arbitrary vertex of another connected component of the graph.
It is convenient to use a queue (note the difference from depth-first search!)
to trace the operation of breadth-first search. The queue is initialized with the
traversal's starting vertex, which is marked as visited. On each iteration, the
algorithm identifies all unvisited vertices that are adjacent to the front vertex,
marks them as visited, and adds them to the queue; after that, the front vertex is
removed from the queue.
Similarly to a DFS traversal, it is useful to accompany a BFS traversal by constructing the so-called breadth-first search forest. The traversal's starting vertex
serves as the root of the first tree in such a forest. Whenever a new unvisited vertex
is reached for the first time, the vertex is attached as a child to the vertex it is being
reached from with an edge called a tree edge. If an edge leading to a previously
visited vertex other than its immediate predecessor (i.e., its parent in the tree) is
encountered, the edge is noted as a cross edge. Figure 5.6 provides an example
of a breadth-first search traversal, with the traversal's queue and corresponding
breadth-first search forest shown.
Here is a pseudocode of the breadth-first search.
ALGORITHM BFS(G)
//Implements a breadth-first search traversal of a given graph
//Input: Graph G = (V, E)
//Output: Graph G with its vertices marked with consecutive integers
//in the order they have been visited by the BFS traversal
5.2
Depth-First Search and Breadth-First Search
169
mark each vertex in V with 0 as a mark of being "unvisited"
count +- 0
for each vertex v in V do
if v is marked with 0
bfs(v)
bfs(v)
//visits all the unvisited vertices connected to vertex v by a path
//and assigns them the numbers in the order they are visited
//via global variable count
count +-count + 1; mark v with count and initialize a queue with v
while the queue is not empty do
for each vertex w in V adjacent to the front vertex do
if w is marked with 0
count +-- count + 1; mark w with count
add w to the queue
remove the front vertex from the queue
Breadth-first search has the same efficiency as depth-first search: it is in
8(1VI 2 ) for the adjacency matrix representation and in 8(1VI +lEI) for the adjacency list representation. Unlike depth-first search, it yields a single ordering of
vertices because the queue is a FIFO (first-in first-out) structure and hence the
order in which vertices are added to the queue is the same order in which they
are removed from it. As to the structure of a BFS forest of an undirected graph,
it can also have two kinds of edges: tree edges and cross edges. Tree edges are the
ones used to reach previously unvisited vertices. Cross edges connect vertices to
those visited before, but, unlike back edges in a DFS tree, they connect vertices
either on the same or adjacent levels of a BFS tree.
Finally, BFS can be used to check connectivity and acyclicity of a graph,
essentially in the same manner as DFS can. It is not applicable, however, for
several less straightforward applications such as finding articulation points. On the
other hand, it can be helpful in some situations where DFS cannot. For example,
BFS can be used for finding a path with the fewest number of edges between two
given vertices. We start a BFS traversal at one of the two vertices given and stop
it as soon as the other vertex is reached. The simple path from the root of the BFS
tree to the second vertex is the path sought. For example, path a-b-e-g in the
graph in Figure 5.7 has the fewest number of edges among all the paths between
vertices a and g. Although the correctness of this application appears to stem
immediately from the way BFS operates, a mathematical proof of its validity is
not quite elementary (see, e.g., [CarOl]).
Table 5.1 summarizes the main facts about depth-first search and breadth-first
search.
r
.
.
170
Decrease-and-Conquer
a
e
b
c
d
f
g
(b)
(a)
FIGURE 5.7 Illustration of the BFS-based algorithm for finding a minimum-edge path.
(a) Graph. (b) Part of its BFS tree that identifies the minimum-edge path
from a to g.
TABLE 5.1 Main facts about depth-first search IDFS) and breadth-first search IBFS)
Data structure
No. of vertex orderings
Edge types (undirected graphs)
Applications
Efficiency for adjacent matrix
Efficiency for adjacent lists
DFS
BFS
stack
2 orderings
tree and back
edges
connectivity,
acyclicity,
articulation
queue
1 ordering
tree and cross
edges
connectivity,
acyclicity,
minimum-edge
points
paths
2
El(IV 11
E>IIVI +IE II
E>(IV 2 11
E>IIVI +)E))
-----Exercises 5 . 2 - - - - - - - - - - - - - - - 1. Consider the following graph.
5.2 Depth-First Search and Breadth-First Search
171
a. Write down the adjacency matrix and adjacency lists specifying this graph.
(Assume that the matrix rows and columns and vertices in the adjacency
lists follow in the alphabetical order of the vertex labels.)
b. Starting at vertex a and resolving ties by the vertex alphabetical order,
traverse the graph by depth-first search and construct the corresponding
depth-first search tree. Give the order in which the vertices were reached
for the first time (pushed onto the traversal stack) and the order in which
the vertices became dead ends (popped off the stack).
2. If we define sparse graphs as graphs for which lEI E O(IVI), which implementation of DFS will have a better time efficiency for such graphs, the one that
uses the adjacency matrix or the one that uses the adjacency lists?
3. Let G be a graph with n vertices and m edges.
a. True or false: All its DFS forests (for traversals starting at different vertices) will have the same number of trees?
b. True or false: All its DFS forests will have the same number of tree edges
and the same number of back edges?
4. Traverse the graph of Problem 1 by breadth-first search and construct the
corresponding breadth-first search tree. Start the traversal at vertex a and
resolve ties by the vertex alphabetical order.
5. Prove that a cross edge in a BFS tree of an undirected graph can connect
vertices only on either the same level or on two adjacent levels of a BFS tree.
6. a. Explain how one can check a graph's acyclicity by using breadth-first
search.
b. Does either of the two traversals-DFS or BPS-always find a cycle faster
than the other? If you answer yes, indicate which of them is better and
explain why it is the case; if you answer no, give two examples supporting
your answer.
7. Explain how one can identify connected components of a graph by using
a. a depth-first search.
b. a breadth-first search.
8. A graph is said to be bipartite if all its vertices can be partitioned into two
disjoint subsets X and Y so that every edge connects a vertex in X with a vertex
in Y. (One can also say that a graph is bipartite if its vertices can be colored in
two colors so that every edge has its vertices colored in different colors; such
graphs are also called 2-colorable). For example, graph (i) is bipartite while
graph (ii) is not.
P:(
Iii!
I!! r!
172
Decrease-and-Conquer
~
&4
(i)
(ii)
a. Design a DFS-based algorithm for checking whether a graph is bipartite.
b. Design a BFS-based algorithm for checking whether a graph is bipartite.
9. Write a program that, for a given graph, outputs
a. vertices of each connected component;
b. its cycle or a message that the graph is acyclic.
10. One can model a maze by having a vertex for a starting point, a finishing point,
dead ends, and all the points in the maze where more than one path can be
taken, and then connecting the vertices according to the paths in the maze.
a. Construct such a graph for the following maze.
b. Which traversal-DFS or BFS-would you use if you found yourself in a
maze and why?
5.3 Topological Sorting
In this section, we discuss an important problem for directed graphs. Before we
pose this problem though, let us review a few basic facts about directed graphs
themselves. A directed graph, or digraph for short, is a graph with directions
specified for all its edges (Figure 5.8a is an example). The adjacency matrix and
adjacency lists are still two principal means of representing a digraph. There are
only two notable differences between undirected and directed graphs in representing them: (1) the adjacency matrix of a directed graph does not have to be
5.3
Topological Sorting
173
//,/~
i
(a)
(b)
FIGURE 5.8 (a) Digraph. (b) DFS forest of the digraph for the DFS traversal started at a.
symmetric; (2) an edge in a directed graph has just one (not two) corresponding
nodes in the digraph's adjacency lists.
Depth-first search and breadth-first search are principal traversal algorithms
for traversing digraphs, but the structure of corresponding forests can be more
complex. Thus, even for the simple example in Figure 5.8a, the depth-first search
forest (Figure 5.8b) exhibits all four types of edges possible in a DFS forest of
a directed graph: tree edges (ab, he, de), back edges (ba) from vertices to their
ancestors, forward edges (ac) from vertices to their descendants in the tree other
than their children, and cross edges (de), which are none of the aforementioned
types.
Note that a back edge in a DFS forest of a directed graph can connect a vertex
to its parent. Whether or not it is the case, the presence of a back edge indicates
that the digraph has a directed cycle. (A directed cycle in a digraph is a sequence
of three or more of its vertices that starts and ends with the same vertex and in
which every vertex is connected to its inllllediate predecessor by an edge directed
from the predecessor to the successor.) Conversely, if a DFS forest of a digraph
has no back edges, the digraph is a dag, an acronym for directed acyclic graph.
Directions on a graph's edges lead to new questions about the graph that are
either meaningless or trivial for undirected graphs. In this section, we discuss one
such problem. As a motivating example, consider a set of five required courses
{Cl, C2, C3, C4, C5) a part-time student has to take in some degree program. The
courses can be taken in any order as long as the following course prerequisites are
met: Cl and C2 have no prerequisites, C3 requires C1 and C2, C4 requires C3, and
C5 requires C3 and C4. The student can take only one course per term. In which
order should the student take the courses?
The situation can be modeled by a digraph in which vertices represent courses
and directed edges indicate prerequisite requirements (Figure 5.9). In terms of this
digraph, the question is whether we can list its vertices in such an order that for
every edge in the graph, the vertex where the edge starts is listed before the vertex
T
174
Decrease-and-Conquer
C1
C3
C2
FIGURE 5.9 Digraph representing the prerequisite structure of five courses
where the edge ends. (Can you find such an ordering of this digraph's vertices?)
This problem is called topological sorting. It can be posed for an arbitrary digraph, but it is easy to see that the problem cannot have a solution if a digraph
has a directed cycle. Thus, for topological sorting to be possible, a digraph must
be a dag. It turns out that being a dag is not only necessary but also sufficient for
topological sorting to be possible; i.e., if a digraph has no cycles, the topological
sorting problem for it has a solution. Moreover, there are two efficient algorithms
that both verify whether a digraph is a dag and, if it is, produce an ordering of
vertices that solves the topological sorting problem.
The first algorithm is a simple application of depth-flrst search: perform a DFS
traversal and note the order in which vertices become dead ends (i.e., are popped
off the traversal stack). Reversing this order yields a solution to the topological
sorting problem, provided, of course, no back edge has been encountered during
the traversal. If a back edge has been encountered, the digraph is not a dag, and
topological sorting of its vertices is impossible.
Why does the algorithm work? When a vertex v is popped off a DFS stack,
no vertex u with an edge from u to v can be among the vertices popped off before
v. (Otherwise, (u, v) would have been a back edge.) Hence, any such vertex u will
be listed after v in the popped-off order list, and before v in the reversed list.
Figure 5.10 illustrates an application of this algorithm to the digraph in Figure 5.9. Note that in Figure 5.10c, we have drawn the edges of the digraph, and
they all point from left to right as the problem's statement requires. It is a con-
C1
C3
C2
C14
(a)
The popping-off order:
C5 1
C42
C3 3
(b)
C5, C4, C3, C1, C2
The topologically sorted list:
C2 5
C2
C1--+C3->-C4_,.C5
~~
(c)
FIGURE 5.10 (a) Digraph for which the topological sorting problem needs to be solved.
(b) DFS traversal stack with the subscript numbers indicating the poppingoff order. (c) Solution to the problem.
--,
5.3 Topological Sorting
175
C1
delete C1
C3
C2
delete C2
C2
delete C3
y
delete C4
delete C5
~
The solution obtained is C1, C2, C3, C4, C5
FIGURE 5.11 Illustration of the source-removal algorithm for the topological sorting
problem. On each iteration, a vertex with no incoming edges is deleted
from the digraph.
venient way to check visually the correctness of a solution to an instance of the
topological sorting problem.
The second algorithm is based on a direct implementation of the decrease (by
one )-and-conquer technique: repeatedly, identify in a remaining digraph a source,
which is a vertex with no incoming edges, and delete it along with all the edges
outgoing from it. (If there are several sources, break the tie arbitrarily. If there is
none, stop because the problem cannot be solved-see Problem 6a.) The order in
which the vertices are deleted yields a solution to the topological sorting problem.
The application of this algorithm to the same digraph representing the five courses
is given in Figure 5.11.
Note that the solution obtained by the source-removal algoritlun is different
from the one obtained by the DFS-based algorithm. Both of them are correct, of
course; the topological sorting problem may have several alternative solutions.
The tiny size of the example we used might create a wrong impression about
the topological sorting problem. But imagine a large project-e.g., in construction
or research-that involves thousands of interrelated tasks with known prerequisites. The first thing you should do in such a situation is to make sure that the set
of given prereqnisites is not contradictory. The convenient way of doing this is
to solve the topological sorting problem for the project's digraph. Only then can
you start thinking about scheduling your tasks to, say, minimize the total completion time of the project. This would require, of course, other algorithms that
you can find in general books on operations research or in special ones on socalled CPM (Critical Path Method) and PERT (Program Evaluation and Review
Technique) methodologies.
T,
176
Decrease-and-Conquer
------Exercises 5 . 3 - - - - - - - - - - - - - - - - 1. Apply the DFS-based algorithm to solve the topological sorting problem for
the following digraphs.
}--.{d
e
(a)
(b)
2. a. Prove that the topological sorting problem has a solution for a digraph if
and only if it is a dag.
b. For a digraph with n vertices, what is the largest number of distinct solutions
the topological sorting problem can have?
3. a. What is the time efficiency of the DFS-based algorithm for topological
sorting?
b. How can one modify the DFS-based algorithm to avoid reversing the
·
vertex ordering generated by DFS?
4. Can one use the order in which vertices are pushed onto the DFS stack
(instead of the order they are popped off it) to solve the topological sorting
problem?
5. Apply the source-removal algorithm to the digraphs of Prohlem 1.
6. a. Prove that a dag must have at least one source.
b. How would you find a source (or determine that such a vertex does not
exist) in a digraph represented by its adjacency matrix? What is the time
efficiency of this operation?
c. How would you find a source (or determine that such a vertex does not
exist) in a digraph represented by its adjacency lists? What is the time
efficiency of this operation?
7. Can you implement the source-removal algorithm for a digraph represented
by its adjacency lists so that its running time is in O(IVI +lEI)?
8. Implement the two topological sorting algorithms in the language of your
choice. Run an experiment to compare their running times.
9. A digraph is called strongly connected if for any pair of two distinct vertices u
and v there exists a directed path from u to v and a directed path from v to u. In
general, a digraph's vertices can be partitioned into disjoint maximal subsets
of vertices that are mutually accessible via directed paths of the digraph; these
subsets are called strongly connected components. There are two DFS-based
5.4
Algorithms for Generating Combinatorial Objects
177
algorithms for identifying strongly connected components. Here is the simpler
(but somewhat less efficient) one of the two.
Step 1 Do a DFS traversal of the digraph given and number its vertices
in the order that they become dead ends.
Step 2 Reverse the directions of all the edges of the digraph.
Step 3 Do a DFS traversal of the new digraph by starting (and, if necessary, restarting) the traversal at the highest numbered vertex
among still unvisited vertices.
The strongly connected components are exactly the subsets of vertices in each
DFS tree obtained during the last traversal.
a. Apply this algorithm to the following digraph to determine its strongly
connected components.
d
e
b. What is the time efficiency class of this algorithm? Give separate answers
for the adjacency matrix representation and adjacency list representation
of an input graph.
c. How many strongly connected components does a dag have?
10. Celebrity problem A celebrity among a group of n people is a person who
knows nobody but is known by everybody else. The task is to identify a
celebrity by only asking questions to people of the form: "Do you know
him/her?" Design an efficient algorithm to identify a celebrity or determine
that the group has no such person. How many questions does your algorithm
need in the worst case?
5.4 Algorithms for Generating Combinatorial Objects
In this section, we keep our promise to discuss algorithms for generating combinatorial objects. The most important types of combinatorial objects are permutations, combinations, and subsets of a given set. They typically arise in problems
that require a consideration of different choices. We already encountered them in
Chapter 3 when we discussed exhaustive search. Combinatorial objects are studied in a branch of discrete mathematics called combinatorics. Mathematicians, of
1F!
,i''
!
178
Decrease-and-Conquer
course, are primarily interested in different counting formulas; we should be grateful for such formulas because they tell us how many items need to be generated.
(In particular, they warn us that the number of combinatorial objects typically
grows exponentially or even faster as a function of the problem's size.) But our
primary interest here lies in algorithms for generating combinatorial objects, not
just in counting them.
Generating Permutations
We start with permutations. For simplicity, we assume that the underlying set
whose elements need to be permuted is simply the set of integers from 1 to n;
more generally, they can be interpreted as indices of elements in ann-element set
{a1 , ... , anl· What would the decrease-by-one technique suggest for the problem
of generating all n! permutations of {1, ... , n)? The smaller-by-one problem is to
generate all (n - 1)! permutations. Assuming that the smaller problem is solved,
we can get a solution to the larger one by inserting n in each of the n possible
positions among elements of every permutation of n - 1 elements. All the permutations obtained in this fashion will be distinct (why?), and their total number will
be n(n- 1)! = n!. Hence, we will obtain all the permutations of {1, ... , n).
We can insert n in the previously generated permutations either left to right
or right to left. It turns out that it is beneficial to start with inserting n into
12 ... (n - 1) by moving right to left and then switch direction every time a new
permutation of {1, ... , n - 1] needs to be processed. An example of applying this
approach bottom up for n = 3 is given in Figure 5.12.
The advantage of this order stems from the fact that it satisfies the minimalchange requirement: each permutation can be obtained from its immediate predecessor by excha1;1ging just two elements in it. (Check this for the permutations
generated in Figure 5.12.) The minimal-change requirement is beneficial both for
the algorithm's speed and for applications using the permutations. For example,
in Section 3.4, we needed permutations of cities to solve the traveling salesman
problem by exhaustive search. If such permutations are generated by a minimalchange algorithm, we can compute the length of a new tour from the length of its
predecessor in constant rather than linear time (how?).
start
insert 2 into 1 right to left
insert 3 into 21 right to left
insert 3 into 21left to right
1
12
123
321
21
132
312
231
213
FIGURE 5.12 Generating permutations bottom up
5.4
Algorithms for Generating Combinatorial Objects
179
It is possible to get the same ordering of permutations of n elements without
explicitly generating permutations for smaller values of n. It can be done by
associating a direction with each element k in a permutation. We indicate such
a direction by a small arrow written above the element in question, e.g.,
---+~----++---
3 2 4 1.
The element k is said to be mobile in such an arrow-marked permutation if its
arrow points to a smaller number adjacent to it. For example, for the permutation
---+ +-----++--3 2 4 1 , 3 and 4 are mobile while 2 and 1 are not. Using the notion of a mobile
element, we can give the following description of the Johnson- Trotter algorithm
for generating permutations.
ALGORITHM
.TohnsonTrotter(n)
//Implements Johnson-Trotter algorithm for generating permutations
//Input: A positive integer n
//Output: A list of all permutations of {1, ... , n)
~~
~
initialize the first permutation with 1 2 ... n
while the last permutation has a mobile element do
find its largest mobile element k ,
swap k and the adjacent integer k.'s arrow points to
reverse the direction of all the elements that are larger thank
add the new permutation to the list
Here is an application of this algorithm for n = 3, with the largest mobile
integer shown in bold:
~~~
123
~~~
132
--312
--~
321
--- --231
213.
This algorithm is one of the most efficient for generating permutations; it can
be implemented to run in time proportional to the number of permutations, i.e.,
in E>(n!). Of course, it is horribly slow for all but very small values of n; however,
this is not the algorithm's fault but rather the "fault" of the problem: it simply asks
to generate too many items.
One can argue that the permutation ordering generated by the JohnsonTrotter algorithm is not quite natural; e.g., the natural place for permutation n
n - 1 ... 1 seems to be the last one on the list. This would be the case if permutations were listed in increasing order-also called the lexicographic order-which
is the order in which they would be listed in a dictionary if the numbers were
interpreted as letters of an alphabet:
123
132 213
231
312 321.
f
180
!
Decrease-and-Conquer
So how can we generate the permutation following a 1a 2 ... a 11 _ 1a 11 in lexicographic order? If a 11 _ 1 < a 11' we can simply transpose these last two elements. For
example, 123 is followed by 132. If a,_ 1 >a,, we have to engage a,_ 2 . If a,_ 2 <
a,_ 1, we should rearrange the last three elements by increasing the (n- 2)th element as little as possible by putting there the next larger than a,_ 2 element chosen
from a,_ 1 and a, and filling positions n - 1 and n with the remaining two of the
three elements a 11 _ 2 , a 11 _ 1, and a11 in increasing order. For example, 132 is followed
by 213 while 231 is followed by 312. In general, we sean a current permutation from
right to left looking for the first pair of consecutive elements a; and a;+ 1 such that
ai < ai+l (and, hence, ai+l > ... > a 11 ). Then we find the smallest element in the
tail that is larger than a;, i.e., min{ajl aj >a;, j > i}, and put it in position i; the
positions from i + 1 through n are filled with the elements a;, a;+l• ... , a,, from
which the element put in the ith position has been eliminated, in increasing order.
For example, 163542 would be followed by 164235. We leave writing a complete
pseudocode of this algorithm for the exercises.
Generating Subsets
Recall that in Section 3.4 we examined the knapsack problem that asks to find
the most valuable subset of items that fits a knapsack of a given capacity. The
exhaustive-search approach to solving this problem discussed there was based on
generating all subsets of a given set of items. In this section, we discuss algorithms
for generating all2" subsets of an abstract set A= {a 1, ... , a,). (Mathematicians
call the set of all subsets of a set its power set.)
The decrease-by-one idea is immediately applicable to this problem, too. All
subsets of A= {a1, ... , a,) can be divided into two groups: those that do not
contain a, and those that do. The former group is nothing but all the subsets of
{a1, ... , a,_ 1}, while each and every element of the latter can be obtained by
adding a, to a subset of {a1, ... , a,_ 1). Thus, once we have a list of all subsets of
{a1, •.. , a,_Jl, we can get all the subsets of {a1, ... , a,) by adding to the list all
its elements with a, put into each of them. An application of this algorithm to
generate all subsets of {a 1, a2 , a3 ) is illustrated in Figure 5.13.
subsets
n
0
0
1
2
3
0
{aJ)
0
{aJ)
{a,)
{a 1, a 2 }
0
{all
{a,)
{a 1, a 2 }
{a,)
{a 1, a 3 }
FIGURE 5.13 Generating subsets bottom up
{a2> a 3 }
{a 1, a2 , a 3 }
5.4 Algorithms for Generating Combinatorial Objects
181
Similarly to generating permutations. we do not have to generate power sets of
smaller sets. A convenient way of solving the problem directly is based on a one-toone correspondence between all211 subsets of ann element set A= {a 1, ... , a11 }
and all 2" bit strings b1 , ... , b, of length n. The easiest way to establish such a
correspondence is to assign to a subset the bit string in which b, = l if belongs
to the subset and b, = 0 if does not belong to it. (We mentioned this idea of
bit vectors in Section 1.4.) For example, the bit string 000 will correspond to the
empty subset of a three-element set, 111 will correspond to the set itself, i.e.,
{a 1, a 2 , a 3 ), while 110 will represent {at, a 2 ). With this correspondence in place,
we can generate all the bit strings of length n by generating successive binary
numbers from 0 to 2" - 1, padded, when necessary, with an appropriate number
of leading O's. For example, for the case of n = 3, we obtain
a,
a,
bit strings
subsets
000
0
001
{a3]
010
011
100
{a2J
{a2, a3)
{atl
101
{at, a3)
Note that while the bit strings are generated by this algorithm in the lexicographic order (in the two-symbol alphabet of 0 and 1), the order of the subsets
looks anything but natural. For example, we might want the so-called squashed
order, in which any subset involving aj can be listed only after all the subsets
involving a 1, ... , aj-l• as was the case for the list of the three-element set in Figure 5.13. It is easy to adjust the bit string-based algorithm to yield a squashed
ordering of the subsets involved (Problem 6).
A more challenging question is whether there exists a minimal-change algorithm for generating bit strings so that every one of them differs from its innnediate
predecessor by only a single bit. (In the language of subsets, we want every subset
to differ from its immediate predecessor by either an addition or a deletion, but
not both, of a single element). The answer to this question is yes (Problem 9); for
example, for n = 3, we can get
000 001
011
010
110 111
101
100.
Such a sequence of bit strings is called the binary reflected Gray code. Gray codes
have many interesting properties and a few useful applications; you can read about
them in such books as [Bru04].
-----Exercises !5.4 - - - - - - - - - - - - - - - 1. Is it realistic to implement an algorithm that requires generating all permutations of a 25-element set on your computer? What about all the subsets of
such a set?
2. Generate all permutations of {1, 2, 3, 4] by
a. the bottom-up minimal-change algorithm.
182
Decrease-and-Conquer
b. the Johnson-Trotter algorithm.
c. the lexicographic-order algorithm.
3. Write a computer program for generating permutations in lexicographic
order.
4. Consider the following implementation of the algorithm for generating permutations discovered by B. Heap [Hea63].
ALGORITHM HeapPermute(n)
//Implements Heap's algorithm for generating permutations
//Input: A positive integer n and a global array A[l..n]
//Output: All permutations of elements of A
ifn=1
write A
else
for i <-- 1 to n do
HeapPermute(n - 1)
if n is odd
swap A[1] and A[n]
else swap A[i] and A[n]
a. Trace the algorithm by hand for n = 2, 3, and 4.
b. Prove correctness of Heap's algorithm.
c. What is the time efficiency of Heap Permute?
5. Generate all the subsets of a four-element set A = (a~o a 2 , a3 , a4 ) by each of
the two algorithms outlined in this section.
6. What simple trick would make the bit string-based algorithm generate subsets
in squashed order?
7. Write a pseudocode for a recursive algorithm for generating all2n bit strings
oflength n.
8. Write a nonrecursive algorithm for generating 2n bit strings of length n that
implements bit strings as arrays and does not use binary additions.
9. a. Use the decrease-by-one technique to generate a Gray code for n = 4.
b. Design a general decrease-by-one algorithm for generating a Gray code of
order n.
10. Design a decrease-and-conquer algorithm for generating all combinations of
k items chosen from n, i.e., all k-element subsets of a given n-element set. Is
your algorithm a minimal-change algorithm?
5.5
Decrease-by-a-Constant-Factor Algorithms
183
11. Gray code and the Tower of Hanoi
a. Show that the disk moves made in the classic recursive algorithm for the
Tower of Hanoi puzzle can be used for generating the binary reflected Gray
code.
b. Show how the binary reflected Gray code can be used for solving the Tower
of Hanoi puzzle.
5.5 Decrease-by-a-Constant-Factor Algorithms
You may recall from the introduction to this chapter that decrease-by-a-constantfactor is the second major variety of decrease-and-conquer. You have already
encountered examples of this design technique in this book: binary search (Section
4.3) and exponentiation by squaring (introduction to Section 5.1 ). In this section,
you will find a few other examples of algorithms based on the decrease-by-aconstant-factor idea. We should not expect a wealth of examples of this kind,
however, because these algorithms are usually logarithmic and, being very fast,
do not happen often; a reduction by a factor other than two is especially rare.
Fake-Coin Problem
Of several versions of the fake-coin identification problem, we consider the one
that best illustrates the decrease-by-a-constant-factor strategy. Among n identically looking coins, one is fake. With a balance scale, we can compare any two sets
of coins. That is, by tipping to the left, to the right, or staying even, the balance
scale will tell whether the sets weigh the same or which of the sets is heavier than
the other but not by how much. The problem is to design an efficient algorithm
for detecting the fake coin. An easier version of the problem~the one we discuss
here~assumes that it is known whether the fake coin is lighter or heavier than
the genuine one 2 (We assume that the fake coin is lighter.)
The most natural idea for solving this problem is to divide n coins into two
piles of Ln/2J coins each, leaving one extra coin apart if n is odd, and put the two
piles on the scale. If the piles weigh the same, the coin put aside must be fake;
otherwise, we can proceed in the same manner with the lighter pile, which must
be the one with the fake coin. Note that even though we divide the coins into
two subsets, after one weighing we are left to solve a single problem of half the
2.
A much more challenging version assumes no additional information about the relative weights of the
fake and genuine coins or even the presence of the fake coin among n given coins. We pursue this more
difficult version in the exercises to Section 11.2.
184
Decrease-and-Conquer
original size. Therefore, according to our classification of the design techniques, it
is a decrease (by half)-and-conquer rather than a divide-and-conquer algorithm.
We can easily set up a recurrence relation for the number of weighings W (n)
needed by this algorithm in the worst case:
W(n) = W(ln/2J) + 1
for n > 1,
W(1) = 0.
This recurrence should look familiar to you. Indeed, it is almost identical to the one
for the worst-case number of comparisons in binary search. (The difference is in
the initial condition.) This similarity is not really surprising, since both algorithms
are based on the same technique of halving an instance size. The solution to the
recurrence for the number of weighings is also very similar to the one we had for
binary search: W(n) = llog2 nj.
This stuff should look elementary by now, if not outright boring. But wait:
the interesting point here is the fact that this algorithm is not the most efficient
solution. We would be better off dividing the coins not into two but into three
piles of about n/3 coins each. (Details of a precise formulation are developed in
the exercises. Do not miss it! If your instructor forgets, demand the instructor to
assign Problem 3.) After weighing two of the piles, we can reduce the instance size
by a factor of three. Accordingly, we should expect the number of weighings to be
about log3 n, which is smaller than log2 n. (Can you tell by what factor?)
Multiplication
a Ia Russe
Now we consider a nonorthodox algorithm for multiplying two positive integers
called multiplication Ia russe, or the Russian peasant method. Let n and m
be positive integers whose product we want to compute, and let us measure the
instance size by the value of n. Now, if n is even, an instance of half the size has
to deal with n/2, and we have an obvious formula relating the solution to the
problem's larger instance to the solution to the smaller one:
a
n
n ·m =-·2m.
2
If n is odd, we need only a slight adjustment of this formula:
n-1
n·m=--·2m+m.
2
Using these formulas and the trivial case of 1· m = m to stop, we can compute
product n · m either recursively or iteratively. An example of computing 50 · 65
with this algorithm is given in Figure 5.14. Note that all the extra addends shown
in parentheses in Figure 5.14a are in the rows that have odd values in the first
column. Therefore we can find the product by simply adding all the elements in
them column that have an odd number in then column (Figure 5.14b).
Also note that the algorithm involves just the simple operations of halving,
doubling, and adding-a feature that might be attractive, for example, to those
5.5
185
Decrease-by-a-Constant-Factor Algorithms
I!
m
11
m
50
25
12
6
3
1
65
130
260
520
1, 040
2,080
( +1040)
50
25
12
6
3
1
65
130
260
520
1,040
2,080
2,080
+(130 + 1040) = 3, 250
(+130)
(a)
130
1, 040
2, 080
3, 250
---------(b)
FIGURE 5.14 Computing 50. 65 by multiplication a Ia russe
who do not want to memorize the table of multiplications. It is this feature of
the algorithm that most prohably made it attractive to Russian peasants who,
according to Western visitors, used it widely in the nineteenth century, and for
whom the method is named. (In fact, the algorithm's idea was used by Egyptian
mathematicians as early as 1650 B.C. [Cha98], p. 16.) It also leads to very fast
hardware implementation since doubling and halving of binary numbers can be
performed using shifts, which are among tbe most basic operations at the machine
level.
Josephus Problem
Our last example is the Josephus problem, named for Flavius Josephus, a famous
Jewish historian who participated in and chronicled the Jewish revolt of 66-70
c.E. against the Romans. Josephus, as a general, managed to hold the fortress of
Jotapata for 47 days, but after the fall of the city he took refuge with 40 diehards in
a nearby cave. There, the rebels voted to perish rather than surrender. Josephus
proposed that each man in turn should dispatch his neighbor, the order to be
determined by casting lots. Josephus contrived to draw the last lot, and, as one
of the two surviving men in the cave, he prevailed upon his intended victim to
surrender to the Romans.
So let n people numbered 1 to n stand in a circle. Starting the grim count
with person number 1., we eliminate every second person until only one survivor
is left. The problem is to determine the survivor's number J(n). For example
(Figure 5 .15), ifn is 6, people in positions2, 4, and 6 will be eliminated on the first
pass through the circle, and people in initial positions 3 and 1 will be eliminated
on the second pass, leaving a sole survivor in initial position 5-thus, I (6) = 5. To
give another example, if n is 7, people in positions 2, 4, 6, and 1 will be eliminated
186
Decrease-and-Conquer
1,
1,
6,
5
4,
2,
7
2,
3,
6,
3,
5,
4,
(b)
(a)
FIGURE 5.15 Instances of the Josephus problem for (a) n = 6 and (b) n = 7. Subscript
numbers indicate the pass on which the person in that position is
eliminated. The solutions are J(6) = 5 and 1(7) = 7, respectively.
on the first pass (it is more convenient to include 1 in the first pass) and people in
positions 5 and, for convenience, 3 on the second-thus, J (7) = 7.
It is convenient to consider the cases of even and odd n's separately. If n is
even, i.e., n = 2k, the first pass through the circle yields an instance of exactly the
same problem but half its initial size. The only difference is in position numbering;
for example, a person in initial position 3 will be in position 2 for the second pass,
a person in initial position 5 will be in position 3, and so on (check Figure 5.15a). It
is easy to see that to get the initial position of a person, we simply need to multiply
his new position by two and subtract one. This relationship will hold, in particular,
for the survivor, i.e.,
J(2k) = 2J(k) -1.
Let us now consider the case of an odd n (n > 1), i.e., n = 2k + 1. The first pass
eliminates people in all even positions. If we add to this the elimination of the
person in position 1 right after that, we are left with an instance of size k. Here, to
get the initial position that corresponds to the new position numbering, we have
to multiply the new position number by two and add one (check Figure 5.15b).
Thus, for odd values of n, we get
J(2k + l) = 2J(k) + 1.
Can we get a closed-form solution to the two-case recurrence (subject to the
initial condition J(l) = 1)? The answer is yes, though getting it requires more
ingenuity than just applying backward substitutions. In fact, one way to find a
solution is to apply forward substitutions to get, say, the first 15 values of J(n),
discern a pattern, and then prove its general validity by mathematical induction.
We leave the execution of this plan to the exercises; alternatively, you can look it
up in [Gra94], whose exposition of the Josephus problem we have been following.
Interestingly, the most elegant form of the closed-form answer involves the binary
representation of size n: J (n) can be obtained by a one-bit cyclic shift left of n itself!
For example, 1(6) = 1(1102 ) = 101 2 = 5 and J(7) = .1(111 2) =111 2 = 7.
5.5
Decrease-by-a-Constant-Factor Algorithms
187
-----Exercises 5 . 5 - - - - - - - - - - - - - - - 1. Design a decrease-by-half algorithm for computing Llog2 n J and determine its
time efficiency.
2. Consider ternary search-the following algorithm for searching in a sorted
array A[O .. n- 1]. If n = 1, simply compare the search key K with the single
element of the array; otherwise, search recursively by comparing K with
A[Lnf3J], and if K is larger, compare it with A[L2nj3J] to determine in which
third of the array to continue the search.
a. What design technique is this algorithm based on?
b. Set up a recurrence for the number of key comparisons in the worst case.
(You may assume that n = 3' .)
c. Solve the recurrence for n = 3k.
d. Compare this algorithm's efficiency with that of binary search.
3. a. Write a pseudocode for the divide-into-three algorithm for the fake-coin
problem. (Make sure that your algorithm handles properly all values of n,
not only those that are multiples of 3.)
b. Set up a recurrence relation for the number ofweighings in the divide-into·
three algorithm for the fake-coin problem and solve it for n = 3k
c. For large values of n, about how many times faster is this algorithm than
the one based on dividing coins into two piles? (Your answer should not
depend on n.)
4. Apply multiplication a Ia russe to compute 26 · 47.
5. a. From the standpoint of time efficiency, does it matter whether we multiply
n by m or m by n by the multiplication a Ia russe algorithm?
b. What is the efficiency class of multiplication a Ia russe?
6. Write a pseudocode for the multiplication a Ia russe algorithm.
7. Find 1(40)-the solution to the Josephus problem for n = 40.
8. Prove that the solution to the Josephus problem is 1 for every n that is a power
of2.
9. For the Josephus problem,
a. compute J(n) for n = 1, 2, ... , 15.
b. discern a pattern in the solutions for the first fifteen values of n and prove
its general validity.
c. prove the validity of getting J (n) by a one-bit cyclic shift left of the binary
representation of n.
188
Decrease-and-Conquer
Variable-Size-Decrease Algorithms
As mentioned in the introduction to this chapter, in the third principal variety
of decrease-and-conquer, the size reduction pattern varies from one iteration of
the algorithm to another, Euclid's algorithm for computing the greatest common
divisor (Section L1) provides a good example of this kind of algorithm. In this
section, we encounter a few more examples of this variety.
Computing a Median and the Selection Problem
The selection problem is the problem of finding the kth smallest element in a list
of n numbers. This number is called the kth order statistic. Of course, for k = 1
or k = n, we can simply scan the list in question to find the smallest or largest
element, respectively. A more interesting case of this problem is for k = rn/21,
which asks to find an element that is greater than one half of the list's elements
and smaller than the other half. This middle value is called the median, and it is
one of the most important quantities in mathematical statistics. Obviously, we can
find the kth smallest element in a list by sorting the list first and then selecting the
kth element in the output of a sorting algorithm. The time of such an algorithm
is determined by the efficiency of the sorting algorithm used. Thus, with a good
sorting algorithm such as mergesort, the algorithm's efficiency is in 0 (n log n).
You should immediately suspect, however, that sorting the entire list is most
likely overkill since the problem asks not to order the entire list but just to find
its kth smallest element. Fortunately, we do have a very efficient (on average)
algorithm for doing a similar task of partitioning an array's elements into two
subsets: the one containing the elements that are less than or equal to some value
p pivoting the partition and the other containing the elements that are greater
than or equal to p:
ait ... ais-1
'----v---'
p ais+l ... ai,.
"S.p
?:.P
~
Such partitioning was the principal part of quicksort, discussed in Chapter 4.
How can we take advantage of a list's partition? Lets be the partition's split
position. If s = k, the pivot p obviously solves the selection problem. (Had we
indexed the list starting at 0, it would have been s = k - 1, of course.) If s > k, the
kth smallest element in the entire list can be found as the kth smallest element in
the left part of the partitioned array. And if s < k, we can proceed by searching
for the (k - s )th smallest element in its right part. Thus, if we do not solve the
problem outright, we reduce its instance to a smaller one, which can be solved by
the same approach, i.e., recursively. In fact, the same idea can be implemented
without recursion as well. For the nonrecursive version, we need not even adjust
the value of k but just continue until s = k.
5.6
Variable-Size-Decrease Algorithms
189
EXAMPLE Find the median of the following list of nine numbers: 4, 1, 10, 9, 7,
12, 8, 2, 15. Here, k = [9/21 = 5 and our task is to find the fifth smallest element
in the array. As earlier, we assume for the sake of convenience that the elements
of the list are indexed from 1 to 9.
We use the same version of array partitioning that we used in our discussion
of quicksort in Chapter 4, which selects the first element as a pivot and rearranges
elements by two oppositely directed scans of the array:
4 1 10 9 7 12 8 2 15
2 1 4 9 7 12 8 10 15
Since s = 3 < k = 5, we proceed with the right part of the list:
9 7 12
8 7 9
8 10 15
12 10 15
Since s = 6 > k = 5, we continue with the left part of the previous sublist:
8 7
7 8
Now s = k = 5, and hence we can stop: the found median is 8, which is greater
than 2, 1, 4, and 7 but smaller than 9, 12, 10, and 15.
Ill
How efficient is this algorithm? We should expect it to be more efficient than
quicksort in the average case because it has to deal with just a single subarray
after a partition while quicksort has to work on two of them. Had the splits always
happened in the middle of a remaining array, the recurrence for the number of
comparisons would have been
C(n) = C(n/2) + (n + 1),
whose solution, according to the Master Theorem (see Chapter 4), is in EJ(n).
Although the array's size is actually reduced in an unpredictable fashion from one
iteration of the algorithm to another (with some size reductions less than half and
some larger), a careful mathematical analysis shows the average-case efficiency to
be the same as it would be had the size always been reduced by one half. In other
words, the algorithm turns out to be linear in the average case. In the worst case,
we have the same embarrassing deterioration of efficiency into 8(n 2 ). Though
computer scientists have discovered an algorithm that works in linear time even
in the worst case [Blo73], it is too complicated to be recommended for practical
applications.
Note also that the partitioning-based algorithm solves a somewhat more general problem of identifying the k smallest and n - k largest clements of a given
list, not just the value of its kth smallest element.
190
Decrease-and-Conquer
Interpolation Search
As the next example of a variable-size-decrease algorithm, we consider an algorithm for searching in a sorted array called interpolation search. Unlike binary
search, which always compares a search key with the middle value of a given sorted
array (and hence reduces the problem's instance size by half), interpolation search
takes into account the value of the search key in order to find the array's element
to be compared with the search key. In a sense, the algorithm mimics the way we
search for a name in a telephone book: if we are searching for someone named
Brown, we open the book not in the middle but very close to the beginning, unlike
our action when searching for someone named, say, Smith.
More precisely, on the iteration dealing with the array's portion between the
leftmost element A[l] and the rightmost element A[r ], the algorithm assumes
that the array's values increase linearly, i.e., along the straight line through the
points (1, A[l]) and (r, A[r]). (The accuracy of this assumption can influence the
algorithm's efficiency but not its correctness.) Accordingly, the search key's value
v is compared with the element whose index is computed as (the roundoff of)
the x coordinate of the point on the straight line through the points (1, A[l]) and
(r, A[r]) whose y coordinate is equal to the search value v (Figure 5.16).
Writing a standard equation for the straight line passing though the points
(I, A[l]) and (r, A[r ]), substituting v for y, and solving it for x leads to the following
formula:
x =I+ l (v- A[l])(r -I) J.
A[r ]- A[l]
value
A lr I ------------------------------ ·-- -------
v~--------~----~
Alii -----------L _______;__________L_____:_________..,. index
x
r
FIGURE 5.16 Index computation in interpolation search
(5.4)
5.6
Variable-Size-Decrease Algorithms
191
The logic behind this approach is quite straightforward. We know that the
array values are increasing (more accurately, not decreasing) from A[l] to A[r ], but
we do not know how they do it. Had the array's values increased linearly, which is
the simplest manner possible, the index computed by formula (5.4) would be the
expected location of the array's element with the value equal to v, Of course, if v
is not between A[l] and A[r ], formula (5.4) need not be applied (why?).
After comparing v with A[x ], the algorithm either stops (if they are equal)
or proceeds by searching in the same manner among the elements indexed either
between! andx- 1 orbetweenx + 1 andr, depending on whether A[x] is smaller
or larger than v. Thus, the size of the problem's instance is reduced, but we cannot
tell a priori by how much.
The analysis of the algorithm's efficiency shows that interpolation search uses
fewer than log2 log2 n + 1 key comparisons on the average when searching in a list
of n random keys. This function grows so slowly that the number of comparisons
will be a very small constant for all practically feasible inputs (see Problem 6). But
in the worst case, interpolation search is only linear, which must be considered as a
bad performance (why?). As a final assessment of the worthiness of interpolation
search versus that of binary search, we can point to an opinion by R. Sedgewick
[Sed88] that binary search is probably better for smaller ftles but interpolation
search is worth considering for large files and for applications where comparisons
are particularly expensive or access costs are very high. Note that in Section 12.4,
we discuss a continuous counterpart of interpolation search, which can be seen as
one more example of a variable-size-decrease algorithm.
Searching and Insertion in a Binary Search Tree
As the last example of this section, let us revisit the binary search tree. Recall
that this is a binary tree whose nodes contain elements of a set of orderable items,
one element per node, so that for every node all elements in the left subtree are
smaller and all the elements in the right subtree are greater than the element
in the subtree's root. When we need to search for an element of a given value
(say, v) in such a tree, we do it recursively in the following manner. If the tree is
empty, the search ends in failure. If the tree is not empty, we compare v with the
tree's root K(r). If they match, a desired element is found and the search can be
stopped; if they do not match, we continue with the search in the left subtree of the
rootifv < K(r) and in therightsubtreeifv > K(r). Thus, oneachiterationofthe
algorithm, the problem of searching in a binary search tree is reduced to searching
in a smaller binary search tree. The most sensible measure of size of a search tree
is its height; obviously, the decrease in a tree's height normally changes from one
iteration to another of the binary tree search-thus giving us an excellent example
of a variable-size-decrease algorithm.
In the worst case of the binary tree search, the tree is severely skewed.
This happens, in particular, if a tree is constructed by successive insertions of an
increasing or decreasing sequence of keys (Figure 5.17).
r
.·
192
Decrease-and-Conquer
~
@
B
e '8
(a)
FIGURE 5.17
@'
J9
~
(b)
Binary search trees for (a) an increasing sequence of keys and (b) a
decreasing sequence of keys
Obviously, the search for a 11 _ 1 in such a tree requires n comparisons, making
the worst-case efficiency of the search operation fall into El(n). Fortunately, the
average-case efficiency turns out to be in El(log n). More precisely, the number of
key comparisons needed for a search in a binary search tree built from n random
keys is about 2ln n "' 1.39log2 n. Since the operation of insertion of a new key into
a binary search tree is almost identical to that of searching there, it also exemplifies
the variable-size-decrease technique and has the same efficiency characteristics as
the search operation.
The Game of Nim
There are several well-known games that share the following features. There
are two players, who move in turn. No randomness or hidden information is
permitted: all players know all information about gameplay. A game is impartial:
each player has the same moves available from the same game position. Each of
a finite number of available moves leads to a smaller instance of the same game.
The game ends with a win by one of the players (there are no ties). The winner is
the last player who is able to move.
A prototypical example of such games is Nim. Generally, the game is played
with several piles of chips, but we consider the one-pile version first. Thus, there is
a single pile of n chips. Two players take turns by removing from the pile at least
one and at most m chips; the number of chips taken may vary from one move t6"
another, but both the lower and upper limits stay the same. Who wins the game
by taking the last chip, the player moving first or second, if both players make the
best moves possible?
Let us call an instance of the game a winning position for the player to
move next if that player has a winning strategy, that is, a sequence of moves that
results in a victory no matter what moves the opponent makes. Let us call an
instance of the game a losing position for the player to move next if every move
available for that player leads to a winning position for the opponent. The standard
approach to determining which positions are winning and which are losing is to
5.6 Variable-Size-Decrease Algorithms
193
6
7
8
9
FIGURE 5.18 Figure 5.18 Illustration of one-pile Nim with the maximum number of
chips that may be taken on each move m = 4. The numbers indicate n, the
number of chips in the pile. The losing positions for the player to move
are circled. Only winning moves from the winning positions are shown (in
bold).
investigate small values of n first. It is logical to consider the instance of n = 0 as
a losing one for the player to move next because this player is the first one who
cannot make a move. Any instance with 1 .:=: n ::; m chips is obviously a winning
position for the player to move next (why?). The instance with n = m + 1 chips
is a losing one because taking any allowed number of chips puts the opponent in
a winning position. (See an illustration form= 4 in Figure 5.18.) Any instance
with m + 2 :0: n :0: 2m + 1 chips is a winning position for the player to move next
because there is a move that leaves the opponent with m + 1 chips, which is a losing
position; 2m + 2 = 2(m + 1) chips is the next losing position, and so on. It is not
difficult to see the pattern that can be formally proved by mathematical induction:
an instance with n chips is a winning position for the player to move next if and only
if n is not a multiple of m + 1. The winning strategy is to taken mod (m + 1) chips
on every move; any deviation from this strategy puts the opponent in a winning
position.
One-pile Nim has been known for a very long time. It appeared, in particular,
as the summation game in the first published book on recreational mathematics,
authored by Claude-Gaspar Bachet, a French aristocrat and mathematician, in
1612: a player picks a positive integer less than, say, ten, and then his opponent
and he take turns adding any integer less than ten; the first player to reach 100
exactly is the winner [Dud70].
·
In general, Nim is played with I > 1 piles of chips of sizes n 1, n 2 , ... , n 1 . On
each move, a player can take any available number of chips, including all of them,
from any single pile. The goal is the same-to be the last player able to make a
move. Note that for I = 2, it is easy to figure out who wins this game and how.
Here is a hint: the answer for the game's instances with n1 = n 2 differs from the
answer for those with n 1 ,P n 2 .
A solution to the general case of Nim is quite unexpected because it is based
on the binary representation of the pile sizes. Let b1, b2 , ... , b 1 be the pile sizes
•
r
194
Decrease-and-Conquer
in binary. Compute their binary digital sum, defined as the sum of binary digits
discarding any carry. (In other words, a binary digits; in the sum is 0 if the number
of 1's in the i th position in the addends is even, and it is 1 if the number of 1 's is
odd.) It turns out that an instance ofNim is a winning one (for the player to move
next) if and only if its binary digital sum contains at least one 1; consequrntly,
Nim's instance is a loosing instance if and only if its binary digit surn contains only
zeros. For example, for the commonly played instance with n 1 = 3, n2 = 4, n3 = 5,
the binary digital sum is
011
100
101
010
Since this sum contains a 1, the instance is a winning one for the payer moving
first. To find a winning move from this position, the player needs to change one of
the three bit strings so that the new binary digital sum contains only O's. It is not
difficult to see that the only way to accomplish this is to remove two chips from
the first pile.
This ingenious solution to the game of Nim was discovered by Harvard mathematics professor C. L. Bouton more than 100 years ago. Since then, mathematicians have developed a much more general theory of such games. An excellent
account of this theory, with applications to many specific games, is given in the
monograph by E. R. Berlekamp, J. H. Conway, and R. K. Guy [Ber03].
----~Exercises
5J5 - - - - - - - - - - - - - - - -
1. a. If we measure the size of an instance of the problem of computing the greatest common divisor of m and n by the size of the second parameter n, by
how much can the size decrease after one iteration of Euclid's algorithm?
b. Prove that the size of an instance will always decrease at least by a factor
of 2 after two successive iterations of Euclid's algorithm.
2. a. Apply the partition-based algorithm to find the median of the list of numbers 9, 12, 5, 17, 20.
b. Show that the worst -case efficiency of the partition-based algorithm for the
selection problem is quadratic.
3. a. Write a pseudocode for a nonrecursive implementation of the partitionbased algorithm for the selection problem.
b. Write a pseudocode for a recursive implementation of this algorithm.
4. Derive the formula underlying interpolation search.
5. Give an example of the worst-case input for interpolation search and show
that the algorithm is linear in the worst case.
Summary
195
6. a. Find the smallest value o£ n for which log2 log2 n + 1 is greater than 6.
b. Determine which, i£ any, of the following assertions are true:
i. log log n E o(log n)
ii. log log n E G(log n)
iii. log log n E Q (log n)
7. a. Outline an algorithm for finding the largest key in a binary search tree.
Would you classify your algorithm as a variable-size decrease algorithm?
b. What is the time efficiency class of your algorithm in the worst case?
8. a. Outline an algorithm for deleting a key from a binary search tree. Would
you classify this algorithm as a variable-size decrease algorithm?
b. What is the time efficiency class of your algorithm in the worst case?
9. Misere one-pile Nim Consider the so-called misere version o£ the one-pile
Nim, in which the player taking the last chip looses the game. All the other
conditions of the game remain the same, i.e., the pile contains n chips and on
each move a player takes at least one but no more than m chips. Identify the
winning and loosing positions (for the player to move next) in this game.
10. a. Moldy chocolate Two players take turns by breaking an m-by-n chocolate
bar, which has one spoiled 1-by-1 square. Each break must be a single
straight line cutting all the way across the bar along the boundaries between
the squares. After each break, the player who broke the bar last eats the
piece that does not contain the spoiled square. TI1e player left with the
spoiled square looses the game. Is it better to go first or second in this
game?
b. Write an interactive program to play this game with the computer. Your
program should make a winning move in a winning position and a random
legitimate move in a losing position.
11. Flipping pancakes There are n pancakes all of different sizes that are stacked
on top of each other. You are allowed to slip a flipper under one of the pancakes and flip over the whole sack above the flipper. The purpose is to arrange
pancakes according to their size with the biggest at the bottom. (You can see
a visualization of this puzzles on the Interactive Mathematics Miscellany and
Puzzles site [Bog].) Design an algorithm for solving this puzzle.
SUMMARY
Iiiii
Decrease-and-conquer is a general algorithm design technique, based on
exploiting a relationship between a solution to a given instance of a problem
and a solution to a smaller instance of the same problem. Once such a
relationship is established, it can be exploited either top down (recursively)
or bottom up (without a recursion).
-..,.
196
Decrease-and-Conquer
!till
There are three major variations of decrease-and-conquer:
decrease-by-a-constant, most often by one (e.g., insertion sort);
decrease-by-a-constant-factor, most often by the factor of two (e.g., binary
search);
variable-size-decrease (e.g., Euclid's algorithm).
"
Insertion sort is a direct application of the decrease (by one )-and-conquer
technique to the sorting problem. It is a (.<) (n 2 ) algorithm both in the worst
and average cases, but it is about twice as fast on average than in the worst case.
The algorithm's notable advantage is a good performance on almost-sorted
arrays.
"
Depth-first search (DFS) and breadth-first search (BFS) are two principal
graph traversal algorithms. By representing a graph in a form of a depth-first
or breadth-first search forest, they help in the investigation of many important
properties of the graph. Both algorithms have the same time efficiency:
8(1VI 2 ) for the adjacency matrix representation and EJ(IVI +lEI) for the
adjacency list representation.
"
A digraph is a graph with directions on its edges. The topological sorting
problem asks to list vertices of a digraph in an order such that for every edge
of the digraph, the vertex it starts at is listed before the vertex it points to.
This problem has a solution if and only if a digraph is a dag (directed acyclic
graph), i.e., it has no directed cycles.
"
There are two algorithms for solving the topological sorting problem. The
first one is based on depth-first search; the second is based on the direct
implementation of the decrease-by-one technique.
OJ
Decrease-by-one technique is a natural approach to developing algorithms
for generating elementary combinatorial objects. The most efficient type
of such algorithms are minimal-change algorithms. However, the number
of combinatorial objects grows so fast that even the best algorithms are of
practical interest only for very small instances of such problems.
ill
till
Identifying a fake coin with a balance scale, multiplication a Ia russe, and the
Josephus problem are examples of problems that can be solved by decreaseby-a-constant-factor algorithms. Two other and more important examples are
binary search and exponentiation by squaring.
For some algorithms based on the decrease-and-conquer technique, the size
reduction varies from one iteration of the algorithm to another. Examples
of such variable-size-decrease algorithms include Euclid's algorithm, the
partition-based algorithm for the selection problem, interpolation search,
and searching and insertion in a binary search tree. Nim exemplifies games
that proceed through a series of diminishing instances of the same game.
I
Transform-and-Conquer
That's the secret to life ... replace one worry with another.
~Charles M. Schulz {1922-2000), American cartoonist, the creator of Peanuts
T
his chapter deals with a group of design methods that are based on the idea
of transformation. We call this general technique transform-and-conquer
because these methods work as two-stage procedures. First, in the transformation
stage, the problem's instance is modified to be, for one reason or another, more
amenable to solution. Then, in the second or conquering stage, it is solved.
There are three major variations of this idea that differ by what we transform
a given instance to (Figure 6.1):
!I
!1:!1
"
transformation to a simpler or more convenient instance of the same
problem-we call it instance simplification
transformation to a different representation of the same instance-we call it
representation change
transformation to an instance of a different problem for which an algorithm
is already available-we call it problem reduction
In the first three sections of this chapter, we encounter examples of the
~stance-simplificatio'ij)variety. Section 6.1 deals with the simple but fruitful idea
of presorting. Many questions about lists are easier to answer if the lists are
sorted. Of course, the benefits of a sorted list should more than compensate for
the time spent on sorting; otherwise, we would be better off dealing with an
simpler instance
problem's
instance
_..
or
another representation
or
_..
solution
another problem's instance
FIGURE 6.1 Transform-and-conquer strategy
197
198
Transform-and-Conquer
unsorted list directl~ectiou 6.2 introduces one of the most important algorithms
in applied mathematics: Gaussian elimination. This algorithm solves a system of
linear equations by first transforming it to another system with a special property
that makes finding a solution quite easy/In Section 6.3, the ideas of instance
simplification and representation change are applied to search trees. The results
are AVL trees and multiway balanced search trees; of the latter we consider the
simplest case, 2-3 trees.
Section 6.4 presents heaps and heapsort. Even if you are already familiar
with this important data structure and its application to sorting, you can still
benefit from looking at them in this new light of transform-and-conquer design.
In Section 6.5, we discuss Horner's rule, a remarkable algorithm for evaluating
polynomials. If there were an Algorithm Hall of Fame, Horner's rule would be a
serious candidate for induction based on the algorithm's elegance and efficiency.
We also consider there two algorithms for the exponentiation problem, both based
on the representation change idea.
The chapter concludes with a review of several applications of the third variety
of transform-and-conquer: problem reduction. This variety should be considered
the most radical of the three: one problem is reduced to another, i.e., transformed
into an entirely different problem. This is a very powerful idea, and it is extensively
used in complexity theory (Chapter 11). Its application to designing practical algorithms is not trivial, however. First, we need to identify a new problem into
which the given problem should be transformed. Then we must make sure that
the transformation algorithm followed by the algorithm for solving the new problem is time efficient compared to other algorithmic alternatives. Among several
examples, we discuss an important special case of mathematical modeling, or
expressing a problem in terms of purely mathematical objects such as variables,
functions, and equations.
1 Presorting
Presorting is an old idea in computer science. In fact, interest in sorting algorithms
is due, to a siguificant degree, to the fact that many questions about a list are
easier to answer if the list is sorted/ Obviously, the time efficiency of algorithms
that involve sorting may depend on the efficiency of the sorting algorithm being
used/For the sake of simplicity, we assume throughout this section that lists are
imp(emented as arrays, because some sorting algorithms are easier to implement
for the array representation.
So far, we have discussed three elementary sorting algorithms-selecti.Qn SPlJ,
l:mbble sort, and insertion sort-that are quadratic in the worst> and\average cases,
. and tw~-;dvanced algorithms-fmerges<:Jrt, which is al~ys in e(n log n)/and
qui<;k:sqrt, whose efficiency is also e (n log n) in the average case)but is quadratic in
the,·woist case) Are there faster sorting algorithms? As we have already stated in
"''" ""'"-~
-~,.,......."·-~~-
-
§'"'-' ---~"~.
-"\
6.1
Presorting
199
Section 1.3 (see also Section 11.2), no general comparison-based sorting algorithm
can have a better efficiency than n log n in the worst case, and the same result holds
for average-case efficiency. 1
Following are three examples that illustrate the idea of presorting. More
examples can be found in the exercises to this section.
EXAMPLE 1
Checking element uniqueness in an array
(f'i)
If this element uniqueness problem looks familiar to you, it should; we considered
a brute-force algorithm for the problem in Section 2.3 (see Example 2). The bruteforce algorithm compared pairs of the array's elements until either two equal
elements were found or no more pairs were left. Its worst-case efficiency was in
EJ(n2).
- -Aiternatively, we can sort the array first and then check only its consecutive
elements: if the array has equal elements, a pair of them must be next to each other
and vice versa.
ALGORITHM
PresortElementUniqueness(A[O.. n -1])
//Solves the element uniqueness problem by sorting the array first
//Input: An array A[O .. n- 1] of orderable elements
//Output: Returns "true" if A has no equal elements, "false" otherwise
sort the array A
for i +- 0 to n - 2 do
if A[i] = A [i + 1] return false
return true
The running time of this algorithm is the sum of the time spent on sorting
and the time spent on checking consecutive elements. Since the former requires
at least n log n comparisons and the latter needs no more than n ~ 1 comparisons,
it is the sorting part that will determine the overall efficiency of the algorithm. So,
if we use a quadratic sorting algorithm here, the entire algorithm will not be more
efficient than the brute-force one. But if we use a good sorting algorithm, such
as mergesort, with worst-case efficiency in EJ(nlog n), the worst-case efficiency of
the entire presorting-based algorithm will be also in El(n log n):
T(n) = Tw,., (n) + TwmCn) E El(n log n) + El(n) = EJ(n log n).
1.
ill
Sorting algorithms called radix sorts are linear but in terms of the total number of input bits. These
algorithms work by comparing individual bits or pieces of keys rather than keys in their entirety.
Although the running time of these algorithms is proportional to the number of input bits, they are
still essentially n log n algorithms because the number of bits per key must be at least log 2 n in order
to accommodate n distinct keys of input.
f
I
200
Transform-and-Conquer
EXAMPLE 2 Computing a mode
A mode is a value that occurs most often in a given list of numbers. For example,
for 5, 1, 5, 7, 6, 5, 7, the mode is 5. (If several different values occur most often,
any of them can be considered a mode.) The brute-force approach to computing a
mode would scan the list and compute the frequencies of all its distinct values, then
find the value with the largest frequency. In order to implement this idea, we can
store the values already encountered, along with their frequencies, in a separate
list. On each iteration, the ith element of the original list is compared with the
values already encountered by traversing this auxiliary list. If a matching value is
found, its frequency is incremented; otherwise, the current element is added to
the list of distinct values seen so far with a frequency of 1.
It is not difficult to see that the worst-case input for this algorithm is a list with
no equal elements. For such a list, its ith element is compared with i - 1 elements
of the auxiliary list of distinct values seen so far before being added to the list with
a frequency of 1. As a result, the worst-case number of comparisons made by this
algorithm in creating the frequency list is
~
0-~
,_,
2
C(n) = L..(i - 1) = 0 + 1 + · · · + (n- 1) =
2
E EJ(n ).
The additional n - 1 comparisons needed to find the largest frequency in the auxiliary list do not change the quadratic worst-case efficiency class of the algorithm.
As an alternative, let us first sort the input. Then all equal values will be
adjacent to each other. To compute the mode, all we need to do is to find the
longest run of adjacent equal values in the sorted array.
ALGORITHM PresortMode(A[O .. n - l]l
//Computes the mode of an array by sorting it first
//Input: An array A[O .. n - 1] of orderable elements
//Output: The array's mode
sort the array A
//current run begins at position i
i+--0
modefrequency+- 0 //highest frequency seen so far
while i :<:0 n - 1 do
runlength+- 1; runvalue+- A[i]
while i+runlength:S n- 1 and A[i+runlength] =runvalue
runlength+-runlength+ 1
if runlength > modefrequency
modefrequency+-runlength; modevalue+-runvalue
i +-- i +runlength
return modevalue
6.1
Presorting
201
The analysis here is similar to the analysis of Example 1: the running time of
the algorithm will be dominated by the time spent on sorting since the remainder
of the algorithm takes linear time (why?). Consequently, with ann log n sort, this
method's worst-ease efficiency will be in a better asymptotic class than the worst
flil
case efficiency of the brute-force algorithm.
EXAMPLE 3 (jeaFCf.li;,(; p--;;;bi;;;;{)
Consider the problem of searching for a given value v in a given array of n sortable
items. The brute-force solution here is sequential search (Section 3.1 ), which needs
n comparisons in the worst case. If the array is sorted first, we can then apply binary
search, which requires only Llog2 nj + l comparisons in the worst case. Assuming
the most efficient n log n sort, the total running time of such a searching algorithm
in the worst case will be
T(n) = Tw1(n) + T""'ch(n) = 8(n log n) + C'J(log n) = 8(n log n),
which is inferior to sequential search. The same will also be true for the averagecase efficiency. Of course, if we are to search in the same list more than once, the
time spent on sorting might well be justified. (Problen{0in the exercises asks to
~
estimate the minimum number of searches needed to justify presorting.)
Before we finish our discussion of presorting, we should mention that many,
if not most, geometric algorithms dealing with sets of points use presortiug in one
way or another. Points can be sorted by one of their coordinates or by their distance
from a particular line or by some angle, and so on. For example, presorting was
used in the divide-and-conquer algorithms for the closest-pair problem and for
the convex-hull problem, which were discussed in Section 4.6.
Also, some problems for directed acyclic graphs can be solved more easily
after topologically sorting the digraph in question. The problems of finding the
shortest and longest paths in such digraphs (see the exercises to Section 9.3)
illustrate this point.
~~~~-Exercises
6.1 ~~~~-~~---------
1. Recall that the median of a list of n numbers is defined as its fn/21 smallest
element. (The median is larger than one half the elements and is smaller than
the other half.) Design a presorting-based algorithm for finding the median
and deterntine its efficiency class.
2. Consider the problem of finding the distance between the two closest numbers
in an array of n numbers. (The distance between two numbers x and y is
computed as lx- yl.)
a. Design a presorting-based algorithm for solving this problem and determine its efficiency class.
r
.
I
202
Transform-and-Conquer
b. Compare the efficiency of this algorithm with that of the brute-force algorithm (see Problem 9 in Exercises 1.2).
3. Let A= {a~o ... , a,} and B = {b" ... , b,j be two sets of numbers. Consider
the problem of finding their intersection, i.e., the set C of all the numbers that
are in both A and B.
a. Design a brute-force algorithm for solving this problem and determine its
efficiency class.
b. Design a presorting-based algorithm for solving this problem and determine its efficiency class.
4. Consider the problem of finding the smallest and largest elements in an array
of n numbers.
a. Design a presorting-based algorithm for solving this problem and determine its efficiency class.
b. Compare the efficiency of the three algorithms: (i) the brute-force algorithm, (ii) this presorting-based algorithm, and (iii) the divide-and-conquer
algorithm (see Problem 2 in Exercises 4.1).
5. Show that the average-case efficiency of one-time searching by the algorithm
that consists of the most efficient comparison-based sorting algorithm followed by binary search is inferior to the average-case efficiency of sequential
search.
6. Estimate how many searches will be needed to justify time spent on presorting
an array of 103 elements if sorting is done by mergesort and searching is done
by binary search. (You may assume that all searches are for elements known
to be in the array.) What about an array of 106 elements?
7. To sort or not to sort? Design a reasonably efficient algorithm for solving each
of the following problems and determine its efficiency class.
a. You are given n telephone bills and m checks sent to pay the bills (n :>: m).
Assuming that telephone numbers are written on the checks, find out who
failed to pay. (For simplicity, you may also assume that only one check is
written for a particular bill and that it covers the bill in full.)
b. You have a file of n student records indicating each student's number,
name, home address, and date of birth. Find out the number of students
from each of the 50 U.S. states.
8. Given a set of n :>: 3 points in the Cartesian plane, connect them in a simple
polygon, i.e., a closed path through all the points so that its line segments
(the polygon's edges) do not intersect (except for neighboring edges at their
common vertex). For example,
6.2
Gaussian Elimination
203
a. Does the problem always have a solution? Does it always have a unique
solution?
b. Design a reasonably efficient algorithm for solving this problem and indicate its efficiency class.
9. You have an array of n numbers and a numbers. Find out whether the array
contains two elements whose sum iss. (For example, for the array 5, 9, 1, 3
and s = 6, the answer is yes, but for the same array and s = 7, the answer is
no.) Design an algorithm for this problem with a better than quadratic time
efficiency.
10. You have a list of n open intervals (a 1, b1), ... , (a,, b,) on the real line. (An
open interval (a, b) comprises all the points strictly between its endpoints a
and b, i.e., (a, b)= (xi a< x < b}.) Find the maximal number of these intervals that have a common point. For example, for the intervals (1, 4), (0, 3),
(-1.5, 2), (3.6, 5), this maximal number is 3. Design an algorithm for this problem with a better than quadratic time efficiency.
11. Anagram detection
a. Design an efficient algorithm for finding all sets of anagrams in a large file
such as a dictionary of English words [BenOO]. For example, eat, ate, and
tea belong to one such set.
b. Write a program implementing the algorithm.
6.2 Gaussian Elimination
You are certainly familiar with systems of two linear equations in two unknowns:
a 11x + a 12 y = b1
azlx + azzY = bz.
Recall that unless the coefficients of one equation are proportional to the coefficients of the other, the system has a unique solution. The standard method for
finding this solution is to use either equation to express one of the variables as a
T
I
204
Transform-and-Conquer
function of the other and then substitute the result into the other equation, yielding a linear equation whose solution is then used for finding the value of the second
variable.
In many applications, we need to solve a system of n equations in n unknowns-
+ a12x2 + · · · + a1 X = b1
llztXl + llzzXz + ... + llznXn = bz
a11x1
11
11
-where n is a large number. Theoretically, we can solve such a system by generalizing the substitution method for solving systems of two linear equations (what
general design technique would such a method be based upon?); however, the
resulting algorithm would be extremely cumbersome.
Fortunately, there is a much more elegant algorithm for solving systems of
linear equations called Gaussian elimination? The idea of Gaussian elimination
is to transform a system of n linear equations in n unknowns to an equivalent
system (i.e., a system with the same solution as the original one) with an uppertriangular coefficient matrix, a matrix with all zeros below its main diagonal:
+ GtzXz + · · · + a1nxn = b1
azlxl + azzXz + ... + aznxn = bz
a11x1
a'nn x n =b'n'
In matrix notations, we can write it as
Ax=b ==} A'x=b',
where
A=
[""
GJ2
a?l
a22
anl
anz
" l ["] [~
azn
~~~
' b=
~2
~
, A' =
'
al2
~
a~z
0
0
""] ["']
a;
a2n
111
,b=
bz
.
b~
(We added primes to the matrix elements and right-hand sides of the new system
to stress the point that their values differ from their counterparts in the original
system.)
2.
T11e method is named after Carl Friedrich Gauss (1777-1855), who~like other giants in the history of
mathematics such as Isaac Newton and Leonhard Euler-made numerous fundamental contributions
to both theoretical and computational mathematics.
6.2
Gaussian Elimination
205
Why is the system with the upper-triangular coefficient matrix better than a
system with an arbitrary coefficient matrix? Because we can easily solve the system
with an upper-triangular coefficient matrix by backward substitutions as follows.
First, we can immediately find the value of x, from the last equation; then we can
substitute this value into the next to last equation to get x,_b and so on until we
substitute the known values of the last n - 1 variables into the first equation, from
which we find the value of x1.
So how can we get from a system with an arbitrary coefficient matrix A to
an equivalent system with an upper-triangular coefficient matrix A'? We can do it
through a series of the so-called elementary operations:
"
Ill
Ill
exchanging two equations of the system
replacing an equation with its nonzero multiple
replacing an equation with a sum or difference of this equation and some
multiple of another equation
Since no elementary operation can change a solution to a system, any system
that is obtained through a series of such operations will have the same solution as
the original one. Let us see how we can get to a system with an upper-triangular
matrix. First, we use a 11 as a pivot to make all x 1 coeffi:?ients zeros in the equations
below the first one. Specifically, we replace the 1second equation with the difference
between it and the first equatimi multiplied by a 2 tfa 11 to get an equation with
a zero coefficient for x 1. Doing the same for the third, fourth, and finally nth
equation-with the multiples a3J!a 11 , a 41 !a11, ... , a,tfa 11 of the first equation,
respectively-makes all the coefficients of x 1 below the first equation zero. Then
we get rid of all the coefficients of x 2 by subtracting an appropriate multiple of the
second equation from each of the equations below the second one. Repeating this
elimination for each of the first n - 1 variables ultimately yields a system with an
upper-triangular coefficient matrix.
Before we look at an example of Gaussian elimination, let ns note that we
can operate with just a system's coefficient matrix augmented, as its (n + 1)st
column, with the equations' right-hand side values. In other words, we need to
write explicitly neither the variable names nor the plus and equality signs.
EXAMPLE Solve the system by Gaussian elimination
2x 1 - x 2 + x 3 = 1
4xl + x 2 - x 3 = 5
X[+ x 2 + x 3 = 0.
Fn'
ij
206
Transform-and-Conquer
[1
-1
1
1 -1
1
1
[~
-1
3
3
1
-3
1
~]
_}]
[~
-1
3
0
1
-3
2
j]
2
2
row 2row 3-
i row 1
i row 1
row3- irow2
Now we can obtain the solution by backward substitutions:
x 3 = (-2)/2 = -1, x 2 = (3- (-3)x3)/3=0, andx1 = (1- x3 - (-l)x2)/2 = 1.
1111
Here is a pseudocode for the elimination stage of Gaussian elimination.
ALGORITHM
GaussElif(linati~n(A[l..n, l..n],
b[l..n])
//Applies Gaussian elimirtation to matrix A of a system's coefficients,
//augmented with vector b of the system's right-hand side values
//Input: Matrix A[J..n, 1, .. n] and column-vector b[l..n]
//Output: An equivalent upper-triangular matrix in place of A with the
//corresponding right-hand side values in the (n + 1)st column
fori +-- 1 ton do A[i, n + 1] +-- b[i] //augments the matrix
fori +--lton-1do
for j +-- i + 1 to n do
fork +-- i to n + 1 do
A[j, k] +-- A[j, k]- A[i, k] * A[j, i] / A[i, i]
There are two important observations to make about this pseudocode: First, it
is not always correct: if A[i, i] = 0, we cannot divide by it and hence cannot use the
ith row as a pivot for the ith iteration of the algorithm. In such a case, we should
take advantage of the first elementary operation and exchange the ith row with
some row below it that has a nonzero coefficient in the ith column. (If the system
has a unique solution, which is the normal case for systems under consideration,
such a row must exist.)
Since we have to be prepared for the possibility of row exchanges anyway, we
can take care of another potential difficulty: the possibility that A[i, i] is so small
and consequently the scaling factor A[j, i]/ A[i, i] so large that the new value of
A[j, k] might become distorted by a round-off error caused by a subtraction of two
numbers of greatly different magnitudes.3 To avoid this problem, we can always
3.
We discuss round-off errors in more detail in Section11.4.
6.2
Gaussian Elimination
207
look for a row with the largest absolute value of the coefficient in the ith column,
exchange it with the ith row, and then use the new A[i, i] as the ith iteration's
pivot. This modification, called partial pivoting, guarantees that the magnitude
of the scaling factor will never exceed 1.
The second observation is the fact that the innermost loop is written with a
glaring inefficiency. Can you find it before checking the following pseudocode,
which both incorporates partial pivoting and eliminates this inefficiency?
ALGORITHM BetterGaussElimination(A[l..n, l..n ], b[l..n])
//Implements Gaussian elimination with partial pivoting
//Input: Matrix A[l..n, 1, .. n] and column-vector b[l..n]
//Output: An equivalent upper-triangular matrix in place of A and the
//corresponding right-hand side values in place of the (11 + 1)st column
fori *--1 to 11 do A[i, 11 + 1] *-- b[i] //appends b to A as the last column
for i *-- 1 to 11 - 1 do
pivotrow+- i
for j *-- i + 1 to n do
if IA[j, i]l > JA[pivotrow,i]l pivot row *-- j
for k *-- i to 11 + 1 do
swap(A[i, k], A[pivotrow,k])
for j *-- i + 1 to n do
temp<-- A[j, i] I A[i, i]
for k <-- i to 11 + 1 do
A[j, k] *-- A[j, k]- A[i, k] *temp
Let us find the time efficiency of this algorithm. Its innermost loop consists of
a single line
A[j, k] <-- A[j, k]- A[i, k] *temp
which contains one multiplication and one subtraction. On most computers, multiplication is unquestionably more expensive than addition/snbtraction, and hence
it is multiplication that is usually quoted as the algorithm's basic operation 4 The
standard summation formulas and rules reviewed in Section 2.3 (see also Appendix A) are very helpful in the following derivation:
4.
As we mentioned in Section 2.1, on some computers, multiplication is not necessarily more expensive
than addition/subtraction. For this algorithm, this point is moot since we can simply count the number
of times the innermost loop is executed, which is, of course, exactly the same number as the number
of multiplications and the number of subtractions there.
208
1
Transform-and-Conquer
n-1
C(n) =
11
n-1
n+l
n
n-1
11
L L L 1 = L L (n + 1- i + 1) = L L (n + 2- i)
i=l j=i+l k=i
i=l j=i+l
i=l j=i+l
n-1
11-l
= l:Cn + 2- i)(n- (i + 1) + 1) = l:Cn + 2- i)(n- i)
i=l
i=l
= (n + l)(n - 1) + n(n - 2) + · · . + 3. 1
n-1
11-l
n-1
J~l
j~1
J~l
= LU + 2)j = L i + L 2 j = (n- 1)n(2n- 1) + 2 (n ~ 1)n
=
n(n -1)(2n + 5)
6
1 3
,
3
"'-11" E b(n ).
6
3
Since the second (backward substitution) stage of the Gaussian elimination
is in ('<)(n 2 ), as you will be asked to show in the exercises, the running time is
dominated by the cubic elimination stage, making the entire algorithm cubic as
well.
Theoretically, Gaussian elimination always either yields an exact solution to a
system of linear equations when the system has a unique solution or discovers that
no such solution exists. In the latter case) the system will have either no solutions
or infinitely many of them. In practice, solving systems of significant size on a
computer by this method is not nearly so straightforward as the method would
lead us to believe. The principal difficulty lies in preventing an accumulation of
round-off errors (see Section 11.4). Consult textbooks on numerical analysis that
analyze this and other implementation issues in great detail.
LU Decomposition and Other Applications
Gaussian elimination has an interesting and very useful by-product called LU
decomposition of the coefficient matrix. In fact, modern commercial implementations of Gaussian elimination are based on such a decomposition rather than on
the basic algorithm outlined earlier.
EXAMPLE Let us return to the example at the beginning of this section, where
we applied Gaussian elimination to the matrix
A=[~
-1
1
1
-n
Consider the lower-triangular matrix L made up of 1's on its main diagonal and
the row multiples used in the Gaussian elimination process
L=
u ~]
1
6.2
Gaussian Elimination
209
and the upper-triangular matrix U that was the result of this elimination
U=
2
-1
1]
[
0
3 -3
0
0
.
2
It turns out that the product LU of these matrices is equal to matrix A. (For this
particular pair of L and U, you can verify this fact by direct multiplication, but as
a general proposition, it needs, of course, a proof, which we will omit.)
Therefore, solving the system Ax = b is equivalent to solving the system
LUx= b. The latter system can be solved as follows. Denote y = Ux, then
Ly =b. Solve the system Ly = b first, which is easy to do because L is a lowertriangular matrix; then solve the system U x = y, with the upper-triangular matrix
U, to find x. Thus, for the system at the beginning of this section, we first solve
Ly=b:
Its solution is
1
1
y1 =1, y2 =5-2y 1 =3, y3 =0- -y 1 - -yz=-2.
2
2
Solving U x = y means solving
and the solution is
Note that once we have the LU decomposition of matrix A, we can solve
systems Ax= b with as many right-hand side vectors bas we want to, one at a time.
This is a distinct advantage over the classic Gaussian elimination discussed earlier.
Also note that the LU decomposition does not actually require extra memory
because we can store the nonzero part of U in the upper-triangular part of A
(including the main diagonal) and store the nontrivial part of L below the main
diagonal of A.
Computing a Matrix Inverse
Gaussian elimination is a very useful algorithm that tackles one of the most
important problems of applied mathematics: solving systems of linear equations.
In fact, Gaussian elimination can also be applied to several other problems of
210
Transform-and-Conquer
linear algebra, such as computing a matrix inverse. The inverse of an n-by-n matrix
A is an n-by-n matrix, denoted A -l, such that
Ar 1 =I,
where l is the n-by-n identity matrix (the matrix with all zero elements except
the main diagonal elements, which are all ones). Not every square matrix has
an inverse, but when it exists, the inverse is unique. If a matrix A does not have
an inverse, it is called singular. One can prove that a matrix is singular if and
only if one of its rows is a linear combination (a sum of some multiples) of the
other rows. A convenient way to check whether a matrix is nonsingular is to apply
Gaussian elimination: if it yields an upper-triangular matrix with no zeros on the
main diagonal, the matrix is nonsingular; otherwise, it is singular. So being singular
is a very special situation, and most square matrices do have their inverses.
Theoretically, inverse matrices are very important because they play the role
of reciprocals in matrix algebra, overcoming the absence of the explicit division
operation for matrices. For example, in a complete analogy with a linear equation
in one unknown ax= b whose solution can be written as x = a- 1b (if a is not
zero), we can express a solution to a system of n equations inn unknowns Ax = b
as x =A - 1b (if A is nonsingular) where b is, of course, a vector, not a number.
According to the definition of the inverse matrix for a nonsingular n-by-n
matrix A, to compute it, we need to find n 2 numbers xij• 1 <( i, j <( n, such that
xv,
X2 11
Xn 11
l[
=
?1 01
0 0
We can find the unknowns by solving n systems of linear equations that have the
same coefficient matrix A, the vector of unknowns xj is the jth column of the
inverse, and the right-hand side vector ej is the jth column of the identity matrix
(1<(j<(n):
We can solve these systems by applying Gaussian elimination to matrix A augmented by the n-by-n identity matrix. Better yet, we can use Gaussian elimination to find the LU decomposition of A and then solve the systems LU xj = ej,
j = 1, ... , n, as explained earlier.
Computing a Determinant
Another problem that can be solved by Gaussian elimination is computing a
determinant. The determinant of an n-by-n matrix A, denoted det A or IAI. is
a number whose value can be defined recursively as follows. If n = 1, i.e., if A
consists of a single element a 11 , det A is equal to a 11 ; for n > 1, det A is computed
6.2
211
Gaussian Elimination
by the recursive formula
"
det A= I:Sjalj det Aj,
J=l
where s j is+ 1 if j is odd and -1 if j is even, a 1j is the element in row 1 and column
j, and A j is the (n - 1)-by-(n - 1) matdx obtained from matrix A by deleting its
row 1 and column j.
In particular, for a 2-by-2 matrix, the definition implies a formula that is easy
to remember:
In other words, the determinant of a 2-by-2 matrix is simply equal to the difference
between the products of its diagonal elements.
For a 3-by-3 matrix, we get
det [
~~: ~~~ ~~~ J
a31
a32
= a11 det [ a22
0 32
a33
a23] - al2 det [ azl
°33
a31
an] + a13 det [ azl
a33
a31
azz]
a32
= a, t azza33+al2a23a3l +azla32al3-a31a22au-a21a12a33-a3zanau.
Incidentally, this formula is very handy in a variety of applications. In particular,
we used it twice already in Section 4.6 as a part of the quickhull algorithm.
But what if we need to compute a determinant of a large matrix? (Although
this is a task that is rarely needed in practice, it is worth discussing nevertheless.)
Using the recursive definition can be of little help because it implies computing
the sum of n! terms. Here, Gaussian elimination comes to the rescue again. The
central point is the fact that the determinant of an upper-triangular matrix is
equal to the product of elements on its main diagonal, and it is easy to see how
elementary operations employed by the algorithm influence the determinant's
value. (Basically, it either remains unchanged or changes a sign or is multiplied
by the constant used by the elimination algorithm.) As a result, we can compute
the determinant of an n-by-n matrix in cubic time.
Determinants play an important role in the theory of systems of linear equations. Specifically, a system of n linear equations in n unknowns Ax = b has a
unique solution if and only if the determinant of its coefficient matrix, det A, is
not equal to zero. Moreover, this solution can be found by the formulas called
Cramer's role:
det A 1
XJ
det Aj
det A,
= det A ' · · · 'xj = det A ' · · · 'x, = det A '
T
212
Transform-and-Conquer
where det A1 is the determinant of the matrix obtained by replacing the jth
column of A by the column b. (You are asked to investigate in the exercises
whether using Cramer's rule is a good algorithm for solving systems of linear
equations.)
-----Exercises 6 . 2 - - - - - - - - - - - - - - - 1. Solve the following system by Gaussian elimination.
x 1 +x2 + x 3 =2
2xt + x2 + x3 = 3
XJ- x2 + 3x3 = 8
2. a. Solve the system of the previous question by the LU-decomposition
method.
b. From the standpoint of general algorithm design techniques, how would
you classify the LU-decomposition method?
3. Solve the system of Problem 1 by computing the inverse of its coefficient
matrix and then multiplying it by the right-hand side vector.
~Would it be correct to get the efficiency class of the elimination stage of
Gaussian elimination as follows?
n-1
n
n+l
n-1
C(n)= I : I : I;l= I:<n+2-i)(n-i)
i=l j=i+l k=i
i=l
n-1
2
= L[(n + 2)n- i(2n + 2) + i ]
i=l
n-l
n-1
n-1
= I;Cn + 2)n- L(2n + 2)i + L
i=l
i=l
;2
i=l
Since s1(n) = L;;':;}Cn + 2)n E El(n 3), s2(n) = L;;'.:;{C2n + 2)i E 8(n 3),
and s3 (n) = "L7.:.J i 2 E El(n 3), s1(n)- s2 (n) + s3 (n) E 8(n 3).
5. Write a pseudocode for the back-substitution stage of Gaussian elimination
and show that its running time is in 8(n 2 ).
@ Assuming that division of two numbers takes three times longer than their
multiplication, estimate how much faster BetterGaussElimination is than
GaussElimination. (Of course, you should also assume that a compiler is not
going to eliminate the inefficiency in GaussElimination.)
7. a. Give an example of a system of two linear equations in two unknowns that
has a unique solution and solve it by Gaussian elimination.
l
6.2
Gaussian Elimination
213
b. Give an example of a system of two linear equations in two unknowns that
has no solution and apply Gaussian elimination to it.
c. Give an example of a system of two linear equations in two unknowns that
has infinitely many solutions and apply Gaussian elimination to it.
8. The Gauss-Jordan elimination method differs from Gaussian elimination in
that the elements above the main diagonal of the coefficient matrix are made
zero at the same time and by the same use of a pivot row as the elements below
the main diagonal.
a. Apply the Gauss-Jordan method to the system of Problem 1 of these
exercises.
b. What general design strategy is this algorithm based on?
c. In general, how many multiplications are made by this method while solving a system of n equations inn unknowns? How does this compare with
the number of multiplications made by the Gaussian elimination method
in both its elimination and back-substitution stages?
9. A system Ax = b of n linear equations in n unknowns has a unique solution if
and only if det A # 0. Is it a good idea to check this condition before applying
Gaussian elimination to a system?
10. a. Apply Cramer's rule to solve the system of Problem 1 of these exercises.
b. Estimate how many times longer it will take to solve a system of n linear
equations in n unknowns by Cramer's rule than by Gaussian elimination.
(Assume that all the determinants in Cramer's rule formulas are computed
independently by Gaussian elimination.)
11. Lights out This one-person game is played on an n-by-n board composed
of 1-by-1light panels. Each panel has a switch that can be turned on and off,
thereby toggling the on/off state of this and four vertically and horizontally
adjacent panels. (Of course, toggling a corner square affects a total of three
panels, and toggling a noncorner panel on the board's border affects a total
of four squares.) Given an initial subset of lighted squares, the goal is to turn
all the lights off.
a. Show that an answer can be found by solving a system of linear equations
with 0/1 coefficients and right -hand sides using modulo 2 arithmetic.
b. Use Gaussian elimination to solve the 2-by-2 "all-ones" instance of this
problem, where all the panels of the 2-by-2 board are initially lit.
c. Use Gaussian elimination to solve the 3-by-3 "all-ones" instance of this
problem, where all the panels of the 3-by-3 board arc initially lit.
~
I
214
Transform-and-Conquer
Balanced Search Trees
In Sections 1.4 and 4.4, we discussed the binary search tree-one of the principal data structures for implementing dictionaries. It is a binary tree whose nodes
contain elements of a set of orderable items, one element per node, so that allele-
ments in the left subtree are smaller than the element in the subtree's root, and all
the elements in the right subtree are greater than it. Note that this transformation
from a set to a binary search tree is an example of the representation-change technique. What do we gain by such transformation compared to the straightforward
implementation of a dictionary by, say, an array? We gain in the time efficiency
of searching, insertion, and deletion, which are all in EJ(log n), but only in the average case. In the worst case, these operations are in 8 (n) because the tree can
degenerate into a severely unbalanced one with its height equal to n - 1.
Computer scientists have expended a lot of effort in trying to find a structure
that preserves the good properties of the classical binary search tree-principally,
the logarithmic efficiency of the dictionary operations and having the set's elements sorted-but avoids its worst-case degeneracy. They have come up with two
approaches.
Ill
Ill
The first approach is of the instance-simplification variety: an unbalanced
binary search tree is transformed to a balanced one. Specific implementations
of this idea differ by their definition of balance. An AVL tree requires the
difference between the heights of the left and right subtrees of every node
never exceed 1. A red-black tree tolerates the height of one subtree being
twice as large as the other subtree of the same node. If an insertion or deletion
of a new node creates a tree with a violated balance requirement, the tree is
restructured by one of a family of special transformations called rotations
that restore the balance required. (In this section, we discuss only AVL trees.
Information about other types of binary search trees that utilize the idea of
rebalancing via rotations, including red-black trees and splay trees, can be
found in the references [Sed98], [Wei98].)
The second approach is of the representation-change variety: allow more than
one element in a node of a search tree. Specific cases of such trees are 2-3
trees, 2-3-4 trees, and more general and important B-trees. They differ in the
number of elements admissible in a single node of a search tree, but all are
perfectly balanced. (We discuss the simplest of such trees, the 2-3 tree, in this
section, leaving the discussion of B-trees for Chapter 7.)
AVl Trees
AVL trees were invented in 1962 by two Russian scientists, G. M. Adelson-Velsky
and E. M. Landis [Ade62], after whom this data structure is named.
6.3 Balanced Search Trees
215
2
10
0
5
20
1
-1
4
7
0
12
0
0
2
8
(a)
0
2
0
0
5
20
1
-1
4
7
0
8
(b)
FIGURE 6.2 lal AVL tree. (b) Binary search tree that is not an AVL tree. The number
above each node indicates that node's balance factor.
DEFINITION AnAVL tree is a binary search tree in which the balance factor of
every node, which is defined as the difference between the heights of the node's
left and right subtrees, is either 0 or +1 or -1. (The height of the empty tree is
defined as -1.)
For example, the binary search tree in Figure 6.2a is an AVL tree but the one
in Figure 6.2b is not.
If an insertion of a new node makes an AVL tree unbalanced, we transform
the tree by a rotation. A rotation in an AVL tree is a local transformation of its
subtree rooted at a node whose balance has become either +2 or -2; if there are
several such nodes, we rotate the tree rooted at the unbalanced node that is the
closest to the newly inserted leaf. There are only four types of rotations; in fact,
two of them are mirror images of the other two. In their simplest form, the four
rotations are shown in Figure 6.3.
The first rotation type is called the single right rotation, orR-rotation. (Imagine rotating the edge connecting the root and its left child in the binary tree in
Figure 6.3a to the right.) Figure 6.4 presents the single R-rotation in its most general form. Note that this rotation is performed after a new key is inserted into the
left subtree of the left child of a tree whose root had the balance of+ 1 before the
insertion.
The symmetric single left rotation, or L-rotation, is the mirror image of the
single R -rotation. It is performed after a new key is inserted into the right subtree
of the right child of a tree whose root had the balance of -1 before the insertion.
(You are asked to draw a diagram of the general case of the single £-rotation in
the exercises.)
The second rotation type is called the double left-right rotation (LR-rotation ).
It is, in fact, a combination of two rotations: we perform the L~rotation of the left
FIGURE 6.3 Four rotation types for AVL trees with three nodes. (a) Single R-rotation.
(b) Single L-rotation. (c) Double LR-rotation. (d) Double RL-rotation.
6.3
Balanced Search Trees
217
single A-rotation
c
c
FIGURE 6.4 General form of the R-rotation in the AVL tree. A shaded node is the last
one inserted.
double LR-rotation
g
c
or
FIGURE 6.5 General form of the double LR-rotation in the AVL tree. A shaded node
is the last one inserted. It can be either in the left subtree or in the right
subtree of the root's grandchild.
subtree of root r followed by the R-rotation of the new tree rooted at r (Figure 6.5).
It is performed after a new key is inserted into the right subtree of the left child
of a tree whose root had the balance of+ 1 before the insertion.
The double right-left rotation (RL-rotation) is the mirror image of the double
LR-rotation and is left for the exercises.
Note that the rotations are not trivial transformations, though, fortunately,
they can be done in constant time. Not only should they guarantee that a resulting
tree is p~lanced, but they should also preserve the basic requirements of a binary
search tree. For example, in the initial tree in Figure 6.4, all the keys of subtree
Transform-and-Conquer
218
T are smaller than c, which is smaller than all the keys of subtree T2 , which are
1
smaller than r, which is smaller than all the keys of subtree T3. And the same
relationships among the key values hold, as they must, for the balanced tree after
the rotation.
An example of constructing an AVL tree for a given list of numbers is shown
in Figure 6.6. As you trace the algorithm's operations, keep in mind that if there
are several nodes with the ±2 balance, the rotation is done for the tree rooted at
the unbalanced node that is the closest to the newly inserted leaf.
How efficient are AVL trees? As with any search tree, the critical characteristic is the tree's height. It turns out that it is bounded both above and below
by logarithmic functions. Specifically, the height h of any AVL tree with n nodes
satisfies the inequalities
Llog2 nj :": h < 1.4405log2 (n + 2)- 1.3277.
..,.
(These weird-looking constants are roundoffs of some irrational numbers related
to Fibonacci numbers and the golden ratio-see Section 2.5.)
The inequalities immediately imply that the operations of searching and insertion are 8 (log n) in the worst case. Getting an exact formula for the average height
.of an AVL tree constructed for random lists of keys has proved to be difficult, but it
is known from extensive experiments that it is about l.Ollog2 n + 0.1 except when
n is small ([Knulll], p. 468). Thus, searching in an AVL tree requires, on average,
almost the same number of comparisons as searching in a sorted array by binary
search. The operation of key deletion in an AVL tree is considerably more difficult than insertion, but fortunately it turns out to be in the same efficiency class as
insertion, i.e., logarithmic.
These impressive efficiency characteristics come at a price, however. The
drawbacks of AVL trees are frequent rotations, the need to maintain balances
for the tree's nodes, and overall complexity, especially of the deletion operation.
(These drawbacks have prevented AVL trees from becoming the standard structure
'for implementing dictionaries. At the same time, their underlying idea-that of
rebalancing a binary search tree via rotations-has proved to be very fruitful and
has led to discoveries of other interesting variations of the classical binary search
tree.
2-3 Trees
As we mentioned at the beginning of this section, the second idea of balancing
a search tree is to allow more than one key in the same node. The simplest
implementation of this idea is 2-3 trees, introduced by the U.S. computer scientist
John Hopcroft in 1970 (see [Aho74]). A 2-3 tree is a tree that can have nodes of
two kinds: 2-nodes and 3-nodes. A 2-node contains a single key K and has two
children: the left child serves as the root of a subtree whose keys are less thanK
and the right child serves as the root of a subtree whose keys are greater than K.
(In other words, a 2-node is the same kind of node we have in the classical binary
•~
'Iff""
..
6.3
0
-1
®
~
Balanced Search Trees
-2
0
o ~ n
5
-1
6
® @
2
6
6
0
8
5
219
1
2
0
8
5
0
3
6
0
3
R(5)
--II>
1
3
0
2
0
8
0
5
0
2
2
0
0
8
0
2
-1
6
LR(6)
--II>
0
2
-1
0
8
0
-2
RLI61
0
2
--II>
0
2
7
FIGURE 6.6 Construction of an AVL tree for the list 5, 6, 8, 3, 2, 4, 7 by successive
insertions. The parenthesized number of a rotation's abbreviation indicates
the root of the tree being reorganized.
I
I
--r
11
I
··'
'
220
Transform-and-Conquer
2-node
I
3-node
K
FIGURE 6.7 Two kinds of nodes of a 2-3 tree
search tree.) A 3-node contains two ordered keys K 1 and K 2 (K 1 < K 2 ) and has
three children. The leftmost child serves as the root of a subtree with keys less
than K 1, the middle child serves as the root of a subtree with keys between Kt and
K 2 , and the rightmost child serves as the root of a subtree with keys greater thau
K 2 (Figure 6.7).
The last requirement of the 2-3 tree is that all its leaves must be on the same
level, i.e., a 2-3 tree is always perfectly height-balanced: the length of a path from
the root of the tree to a leaf must be the same for every leaf. It is this property that
we "buy" by allowing more than one key in the same node of a search tree.
Searching for a given key Kin a 2-3 tree is quite straightforward. We start at
the root. If the root is a 2-node, we act as if it were a binary search tree: we either
stop if K is equal to the root's key or continue the search in the left or right subtree
if K is, respectively, smaller or larger than the root's key. If the root is a 3-node, we
know after no more than two key comparisons whether the search can be stopped
(if K is equal to one of the root's keys) or in which of the root's three subtrees it
needs to be continued.
Inserting a new key in a 2-3 tree is done as follows. First of all, we always insert
a new key [( in a leaf, except for the empty tree. The appropriate leaf is found by
performing a search for K. If the leaf in question is a 2-node, we insert K there
as either the tirst or the second key, depending on whether [( is smaller or larger
than the node's old key. If the leaf is a 3-node, we split the leaf in two: the smallest
of the three keys (two old ones and the new key) is put in the first leaf, the largest
key is put in the second leaf, while the middle key is promoted to the old leaf's
parent. (If the leaf happens to be the tree's root, a new root is created to accept
the middle key.) Note that promotion of a middle key to its parent can cause the
parent's overflow (if it was a 3-node) and hence can lead to several node splits
along the chain of the leaf's ancestors.
An example of a 2-3 tree construction is given in Figure 6.8.
As for any search tree, the efficiency of the dictionary operations depends on
the~tre'y's hdghl. So let us tirst find an upper bound for it. A 2-3 tree of height h
with the siT1~llest number of keys is a full tree of 2-nodes (such as the final tree in
6.3 Balanced Search Trees
221
~
0@0
5
3
2
9
2
9
2
4
9
FIGURE 6.8 Construction of a 2-3 tree for the list 9, 5, 8, 3, 2, 4, 7
Figure 6.8 for h = 2). Therefore, for any 2-3 tree of height h with n nodes, we get
the inequality
and hence
On the other hand, a 2-3 tree of height h with the largest number of keys is a full
tree of 3-nodes, each with two keys and three children. Therefore, for any 2-3 tree
with n nodes,
n :<' 2. 1 + 2 · 3 + · · · + 2 · 3" = 2(1 + 3 + .. · + 3") = 3"+1- 1,
and hence
These lower and upper bounds on height h,
imply that the time efficiencies of searching, insertion, and deletion are all in
El(log n) in both the worst and average case. We consider a very important generalization of 2-3 trees, called B-trees, in Section 7.4.
222
Transform-and-Conquer
-----Exercises 6 . 3 - - - - - - - - - - - - - - - 1. Which of the following binary trees are AVL trees?
5
... \
6 \
3
c 2
8
(a)
2. a. For n = 1, 2, 3, 4, and 5, draw all the binary trees with n nodes that satisfy
the balance requirement of AVL trees.
b. Draw a binary tree of height 4 that can be an AVL tree and has the smallest
number of nodes among all such trees.
3, ·.Draw diagrams of the single £-rotation and of the doubleR L-rotation in their
general form.
4. For each of the following lists, construct an AVL tree by inserting their elements successively, starting with the empty tree.
a. 1, 2, 3, 4, 5, 6
b. 6, 5, 4, 3, 2, 1
c. 3, 6, 5, 1, 2, 4
5. a. For an AVL tree containing real numbers, design an algorithm for computing the range (i.e., the difference between the largest and smallest numbers
in the tree) and determine its worst-case efficiency.
b. True or false: The smallest and the largest keys in an AVL tree can always
be found on either the last level or the next-to-last level?
6. Write a program for constructing an AVL tree for a given list of n distinct
integers.
7. a. Constructa2-3 tree for the listC, 0, M,P, U, T,I,N, G. (Use the alphabetical
order of the letters and insert them successively starting with the empty
tree.)
b. Assuming that the probabilities of searching for each of the keys (i.e., the
letters) are the same, find the largest number and the average number of
key comparisons for successful searches in this tree.
6.4
Heaps and Heapsort
223
8. Let Ts and T2_3 be, respectively, a classical binary search tree and a 2-3 tree
constructed for the same list of keys inserted in the corresponding trees in
the same order. True or false: Searching for the same key in T2 _3 always takes
fewer or the same number of key comparisons as searching in Ts?
9. For a 2-3 tree containing real numbers, design an algorithm for computing
the range (i.e., the difference between the largest and smallest numbers in the
tree) and determine its worst-case efficiency.
10. Write a program for constructing a 2-3 tree for a given list of n integers.
6.4 Heaps and Heapsort
The data structure called the "heap" is definitely not a disordered pile of items
as the word's definition in a standard dictionary might suggest. Rather it is a
clever, partially ordered data structure that is especially suitable for implementing
priority queues. Recall that a priority queue is a multiset of items with an orderable
characteristic called an item's priority, with the following operations:
"
"
"
finding an item with the highest (i.e., largest) priority
deleting an item with the highest priority
adding a new item to the multiset
It is primarily an efficient implementation of these operations that makes the
heap both interesting and useful. The heap is also the data structure that serves as
a cornerstone of a theoretically important sorting algorithm called heapsort. We
discuss this algorithm after we define the heap and investigate its basic properties.
Notion of the Heap
DEFINITION A heap can be defined as a binary tree with keys assigned to its
nodes (one key per node) provided the following two conditions are met:
1.
2.
The tree's shape requirement-the binary tree is essentially complete (or
simply complete), that is, all its levels are full except possibly the last level,
where only some rightmost leaves may be missing.
The parental dominance requirement-the key at each node is greater than or
equal to the keys at its children. (This condition is considered automatically
satisfied for all leaves. )5
For example, consider the trees of Figure 6.9. The first tree is a heap. The
second one is not a heap, because the tree's shape requirement is violated. And
5.
Some authors require the key at each node to be less than or equal to the keys at its children. We call
this variation a min-heap.
T
224
I
Transform-and-Conquer
6
4
FIGURE 6.9 Illustration of the definition of "heap": only the leftmost tree is a heap.
the third one is not a heap, because the parental dominance requirement fails for
the node with key 5.
Note that key values in a heap are ordered top down; that is, a sequence of
values on any path from the root to a leaf is decreasing (nonincreasing, if equal
keys are allowed). However, there is no left-to-right order in key values; that is,
there is no relationship among key values for nodes either on the same level of
the tree or, more generally, in the left and right subtrees of the same node.
Here is a list of important properties of heaps, which are not difficult to prove
(check these properties for the heap of Figure 6.10, as an example).
There exists exactly one essentially complete binary tree with n nodes. Its
height is equal to Llog2 nj.
2. The root of a heap always contains its largest element.
3. A node of a heap considered with all its descendants is also a heap.
4. A heap can be implemented as an array by recording its elements in the topdown, left-to-right fashion. It is convenient to store the heap's elements in
positions 1 though n of such an array, leaving H[O] either unused or putting
there a sentinel whose value is greater than every element in the heap. In such
1.
a representation,
a.
the parental node keys will be in the first ln/2J positions of the array,
while the leaf keys will occupy the last rn/21 positions;
b. the children of a key in the array's parental position i (1:::: i:::: Ln/2J) will
be in positions 2i and 2i + 1, and, correspondingly, the parent of a key in
position i (2 :::: i :::: n) will be in position li /2 J.
the array representation
index ,o"--r-:'::1-,-"2-,~3-,:::-4--r5~_c6',
value!L _Ll1~0~~,=.,5t;l7'--Ll4~10::2"::-!'---1I
4
FIGURE 6.10 Heap and its array representation
parents
leaves
6.4 Heaps and Heapsort
225
6
6
2
FIGURE 6.11 Bottom-up construction of a heap for the list 2, 9, 7, 6, 5, 8
Thus, we could also define a heap as an array H[l..n] in which every element
in position i in the first half of the array is greater than or equal to the elements
in positions 2i and 2i + 1, i.e.,
H[i]2' max{H[2i], H[2i + 1]} fori= 1, ... , Ln/2J.
(Of course, if2i + 1 > n, just H[i]2' H[2i] needs to be satisfied.) While the ideas
behind the majority of algorithms dealing with heaps are easier to understand if
we think of heaps as binary trees, their actual implementations are usually much
simpler and more efficient with arrays.
How can we construct a heap for a given list of keys? There are two principal
alternatives for doing this. The first is the bottom-up heap construction algorithm
(illustrated in Figure 6.11). It initializes the essentially complete binary tree with
n nodes by placing keys in the order given and then "heapifies" the tree as
follows. Starting with the last parental node, the algorithm checks whether the
parental dominance holds for the key at this node. If it does not, the algorithm
exchanges the node's key K with the larger key of its children and checks whether
the parental dominance holds for K in its new position. 1bis process continues
until the parental dominance requirement for K is satisfied. (Eventually it has
to because it holds automatically for any key in a leaf.) After completing the
"heapification" of the subtree rooted at the current parental node, the algorithm
proceeds to do the same for the node's immediate predecessor. The algorithm
stops after this is done for the tree's root.
Before we introduce a pseudocode of the bottom-up heap construction algorithm, we should make one more remark. Since the value of a node's key does
not change during the process of sifting it down the tree, there is no need to involve it in intermediate swaps. You can think of this improvement as swapping the
empty node with larger keys in its children until a final position is reached where
it accepts the "erased" value again.
226
Transform-and-Conquer
ALGORITHM HeapBottomUp(H[l..n])
//Constructs a heap from the elements of a given array
II by the bottom-up algorithm
//Input: An array H[l..n] of orderable items
//Output: A heap H[l..n]
fori +- Ln/2J downto 1 do
k +-i; v +-H[k]
heap+- false
while not heap and 2 * k :" n do
j +-2*k
if j < n //there are two children
ifH(j]<H(j+1] j +-j+1
if v :>: H[j]
heap+- true
else H[k] +- H[j]; k +- j
H[k] +- v
How efficient is this algorithm in the worst case? Assume, for simplicity, that
n = 2k - 1 so that a heap's tree is full, i.e., the maximum number of nodes occurs on
each level. Let h be the height of the tree; according to the first property of heaps
in the list at the beginning of the section, h = Llog2 nJ (or just rlog 2 (n + l)l - 1 =
k - 1 for the specific values of n we are considering). Each key on level i of
the tree will travel to the leaf level h in the worst case of the heap construction
algorithm. Since moving to the next level down requires two comparisons-one to
find the larger child and the other to determine whether the exchange is requiredthe total number of key comparisons involving a key on level i will be 2(h - i).
Therefore, the total number of key comparisons in the worst case will be
h-l
Cwm-st(n) =
L
I;
i =0
level i keys
h-1
2(h- i)
=I; 2(h- i)i =2(n -log2 (n + 1)),
where the validity of the last equality can be proved either by using the closed-form
formula for the sum L;'~ 1 i2; (see Appendix A) or by mathematical induction on
h. Thus, with this bottom-up algorithm, a heap of size n can be constructed with
fewer than 2n comparisons.
The alternative (and less efficient) algorithm constructs a heap by successive
insertions of a new key into a previously constructed heap; some people call it
the top-down heap construction algorithm. So how can we insert a new key K
into a heap? First, attach a new node with key K in it after the last leaf of the
existing heap. Then sift K up to its appropriate place in the new heap as follows.
Compare K with its parent's key: if the latter is greater than or equal to K, stop
(the structure is a heap); otherwise, swap these two keys and compare K with its
new parent. This swapping continues until K is not greater than its last parent or
6.4
2
10
2
Heaps and Heapsort
8
2
227
8
FIGURE 6.12 Inserting a key (1 0) into the heap constructed in Figure 6.11. The new key
is sifted up via a swap with its parent until it is not larger than its parent
(or is in the root).
it reaches the root (illustrated in Figure 6.12). In this algorithm, too, we can sift
up an empty node until it reaches its proper position, where it will get K 's value.
Obviously, this insertion operation cannot require more key comparisons than
the heap's height. Since the height of a heap with n nodes is about log2 n, the time
efficiency of insertion is in O(log n).
How can we delete an item from a heap? We consider here only the most
important case of deleting the root's key, leaving the question about deleting an
arbitrary key in a heap for the exercises. (Authors of textbooks like to do such
things to their readers, do they not?) So deleting the root's key from a heap can
be done with the following algorithm (illustrated in Figure 6.13).
Maximum Key Deletion from a heap
Step 1 Exchange the root's key with the last key K of the heap.
Step 2 Decrease the heap's size by 1.
Step 3 "Heapify" the smaller tree by sifting K down the tree exactly in the
same way we did it in the bottom-up heap construction algorithm. That
is, verify the parental dominance for K: if it holds, we are done; if not,
swap K with the larger of its children and repeat this operation until
the parental dominance condition holds forK in its new position.
The efficiency of deletion is determined by the number of key comparisons
needed to "heapify" the tree after the swap has been made and the size of the tree
is decreased by 1. Since it cannot require more key comparisons than twice the
heap's height, the time efficiency of deletion is in O(log n) as well.
Heapsort
Now we can describe heapsort-an interesting sorting algorithm discovered by
J. W. J. Williams [Wil64]. This is a two-stage algorithm that works as follows.
Stage 1 (heap construction): Construct a heap for a given array.
Stage 2 (maximum deletions): Apply the root-deletion operation n - 1 times
to the remaining heap.
Transform-and-Conquer
228
8
2
Step 2
Step 1
-+
-+
2
5
2
9
6
6
5
2
FIGURE 6.13 Deleting the root's key from a heap. The key to be deleted is swapped
with the last key, after which the smaller tree is "heapified" by exchanging
the new key at its root with the larger key at its children until the parental
dominance requirement is satisfied.
As a result, the array elements are eliminated in decreasing order. But since
under the array implementation of heaps, an element being deleted is placed last,
the resulting array will be exactly the original array sorted in ascending order.
Heapsort is traced on a specific input in Figure 6.14. (The same input as the one
of Figure 6.11 is intentionally used so that you can compare the tree and array
implementations of the bottom-up heap construction algorithm.)
Since we already know that the heap construction stage of the algorithm is in
O(n), we need to investigate just the time efficiency of the second stage. For the
number of key comparisons, C (n), needed for eliminating the root keys from the
heaps of diminishing sizes from n to 2, we get the following inequality:
n-l
C(n) :=: 2Llog2 (n -1)j + 2Liog2 (n- 2)j + · · · + 2Llog2 lj S 2
L log2 i
i=l
n-1
:=: 2 L log2 (n - 1) = 2(n - 1) log2 (n - 1) S 2n log2 n.
i=l
This means that C(n) E O(n log n) for the second stage ofheapsort. For both stages,
we get O(n) + O(n log n) = O(n log n). A more detailed analysis shows that the
time efficiency of heapsort is, in fact, in E>(n log n) in both the worst and average
cases. Thus, heapsort's time efficiency falls in the same class as that of mergesort;
but, unlike the latter, it is in place, i.e., it does not require any extra storage. Timing
experiments on random files show that heapsort runs more slowly than quicksort
but is competitive with mergesort.
6.4 Heaps and Heapsort
Stage 1 (heap construction)
229
Stage 2 (maximum deletions)
2
9
7
6
5
8
9
6
8
2
5
2
9
8
6
5
7
7
6
8
2
5 19
2
9
8
6
5
7
8
6
7
2
5
9
2
8
6
5
7
5
6
7
2
8
9
6
8
2
5
7
7
6
5
2
2
6
517
6
2
5
5
2
6
5
2
7
215
2
FIGURE 6.14 Sorting the array 2, 9, 7, 6, 5, 8 by heapsort
-----Exercises 6 . 4 - - - - - - - - - - - - - - - 1. a. Construct a heap for the list 1, 8, 6, 5, 3, 7, 4 by the bottom-up algorithm.
b. Construct a heap for the list 1, 8, 6, 5, 3, 7, 4 by successive key insertions
(top-down algorithm).
c. Is it always true that the bottom-up and top-down algorithms yield the
same heap for the same input?
2. Outline an algorithm for checking whether an array H[l..n] is a heap and
determine its time efficiency.
3. a. Find the smallest and the largest number of keys that a heap of height h
can contain.
b. Prove that the height of a heap with n nodes is equal to Llog2 n J.
4. Prove the following equality used in Section 6.4
h-1
L 2(h- i)zi = 2(n -log2(n + 1)), where n = zi•+l- 1.
i=O
230
Transform-and-Conquer
5. a. Design an efficient algorithm for finding and deleting an element of the
smallest value in a heap and determine its time efficiency.
b. Design an efficient algorithm for finding and deleting an element of a given
value v in a given heap H and determine its time efficiency.
6. Sort the following lists by heapsort by using the array representation of heaps.
a. 1, 2, 3, 4, 5 (in increasing order)
b. 5, 4, 3, 2, 1 (in increasing order)
c. S, 0, R, T, I, N, G (in alphabetical order)
7. Is heapsort a stable sorting algorithm?
8. What variety of the transform-and-conquer technique does heapsort represent?
9. Which sorting algorithm other than heapsort uses a priority queue?
10. Implement three advanced sorting algorithms-mergesort, quicksort, and
heapsort-in the language of your choice and investigate their performance
on arrays of sizes n = 102 , 103 , 104 , 105 and 106 For each of these sizes,
consider
a. randomly generated files of integers in the range [l..n].
b. increasing files of integers 1, 2, ... , n.
c. decreasing ftles of integers n, n - 1, ... , 1.
11. Spaghetti sort Imagine a handful of uncooked spaghetti, individual rods
whose lengths represent numbers that need to be sorted.
a. Outline a "spaghetti sort"-a sorting algorithm that takes advantage of
this unorthodox representation.
b. What does this example of computer science folklore (see [Dew93]) have
to do with the topic of this chapter in general and heapsort in particular?
6.5 Horner's Rule and Binary Exponentiation
In this section we discuss the problem of computing the value of a polynomial
(6.1)
at a given point x and its important special case of computing xn. Polynomials
constitute the most important class of functions because they possess a wealth of
good properties on the one hand and can be used for approximating other types
of functions on the other. The problem of manipulating polynomials efficiently
has been important for several centuries; new discoveries were still being made
in the last 50 years. By far the most important of them was the fast Fourier
1
6.5
Horner's Rule and Binary Exponentiation
231
transform (FFT). The practical importance of this remarkable algorithm, which
is based on representing a polynomial by its values at specially chosen points,
was such that some people consider it one of the most important algorithmic
discoveries of all times. Because of its relative complexity, we do not discuss the
FFT algorithm in this book. An interested reader will find a wealth of literature
on the subject, including reasonably accessible treatments in such textbooks as
[Sed88] and [CarOl].
Horner's Rule
Homer's rule is an old but very elegant and efficient algorithm for evaluating a
polynomial. It is named after the British mathematician W. G. Horner, who published it in the early 19th century. But according to Knuth ([Knuii], p. 486), the
method was used by Isaac Newton 150 years before Horner. You will appreciate
this method much more if you first design an algorithm for the polynomial evaluation problem by yourself and investigate its efficiency (see Problems 1 and 2 in
the exercises to this section).
Horner's rule is a good example of the representation change technique since
it is based on representing p(x) by a formula different from (6.1). This new formula
is obtained from (6.1) by successively takingx as a common factor in the remaining
polynomials of diminishing degrees:
p(x) = ( ... (a,x + a,_ 1)x + .. .)x + a0 .
(6.2)
For example, for the polynomial p(x) = 2x 4 - x 3 + 3x 2 + x- 5, we get
p(x)=2x 4 -x 3 +3x 2 +x -5
+ 3x + 1) - 5
= x(x(2x - x + 3) + 1)- 5
= x(x(x(2x- 1) + 3) + 1)- 5.
3
= x (2x - x
2
2
(6.3)
It is in formula (6.2) that we will substitute a value of x at which the polynomial needs to be evaluated. It is hard to believe that this is a way to an efficient
algorithm, but the unpleasant appearance of formula (6.2) is just that, an appearance. As we shall see, there is no need to go explicitly through the transformation
leading to it: all we need is an original list of the polynomial's coefficients.
The pen-and-pencil calculation can be conveniently organized with a tworow table. The first row contains the polynomial's coefficients (including all the
coefficients equal to zero, if any) listed from the highest a, to the lowest a0 . Except
for its first entry, which is a,, the second row is filled left to right as follows: the
next entry is computed as the x 's value times the last entry in the second row plus
the next coefficient from the first row. The final entry computed in this fashion is
the value being sought.
232
l
Transform-and-Conquer
EXAMPLE 1 Evaluate p(x) = 2x 4 - x 3 + 3x 2 + x- 5 at x = 3.
coefficients
x=3
2
-1
1
3
-5
2 3-2+(-1)=5 3-5+3=18 3·18+1=55 3-55+(-5)=160
So p(3) = 160. (On comparing the table's entries with formula (6.3), you will see
that 3 . 2 + (-1) = 5 is the value of 2x - 1 at x = 3, 3 . 5 + 3 = 18 is the value of
x(2x- 1) + 3 at x = 3, 3 · 18 + 1 =55 is the value of x(x(2x- 1) + 3) + 1 at x = 3,
and finally, 3. 55+ ( -5) = 160 is the value of x(x(x(2x- 1) + 3) + 1)- 5 = p(x)
at x = 3.)
Ill
The pseudocode for this algorithm is the shortest one imaginable for a nontrivial algorithm:
ALGORITHM Horner(P[O .. n], x)
//Evaluates a polynomial at a given point by Horner's rule
//Input: An array P[O .. n] of coefficients of a polynomial of degree n
II
(stored from the lowest to the highest) and a number x
//Output: The value of the polynomial at x
p <- P[n]
for i <- n - 1 downto 0 do
p<-X*p+P[i]
return p
The number of multiplications and the number of additions are given by the
same sum:
n-1
M(n) = A(n) =
L 1 = n.
i=O
To appreciate how efficient Horner's rule is, consider only the first term of
a polynomial of degree n: anx". Just computing this single term by the bruteforce algorithm would require n multiplications, whereas Horner's rule computes,
in addition to this term, n - 1 other terms, and it still uses the same number of
multiplications! It is not surprising that Horner's rule is an optimal algorithm for
polynomial evaluation without preprocessing the polynomial's coefficients. But it
took scientists 150 years after Horner's publication to come to the realization that
such a question was worth investigating.
Horner's rule also has some useful by-products. The intermediate numbers
generated by the algorithm in the process of evaluating p(x) at some point x 0 turn
6.5
Horner's Rule and Binary Exponentiation
233
out to be the coefficients of the quotient of the division of p(x) by x - x 0 , while the
final result, in addition to being p(x0 ), is equal to the remainder of this division.
Thus, according to Example 1, the quotient and the remainder of the division of
4
3
2
2x - x + 3x + x - 5 by x - 3 are 2x 3 + 5x 2 + 18x +55 and 160, respectively.
This division algorithm, known as synthetic division, is more convenient than
so-called long division. (However, unlike long division, it is applicable only to
divisions by x - c where c is some constant.)
Binary Exponentiation
The amazing efficiency of Horner's rule fades if the method is applied to computing a", which is the value of x 11 at x =a. In fact, it degenerates to the brute-force
multiplication of a by itself, with wasteful additions of zeros in between. Since
computing an (actually, a 11 mod m) is an essential operation in several important
primality-testing and encryption methods, we consider now two algorithms for
computing a" that are based on the representation change idea. They both exploit
the binary representation of exponent n, but one of them processes this binary
string left to right whereas the second does it right to left.
Let
n=b 1 .. . b, .. . b0
be the bit string representing a positive integer n in the binary number system.
This means that the value of n can be computed as the value of the polynomial
I
.
p(x) = b1 x + · · · + h;x' + · · · + bo
(6.4)
at x = 2. For example, if n = 13, its binary representation is 1101 and
Let us now compute the value of this polynomial by applying Horner's rule
and see what the method's operations imply for computing the power
Horner's rule for the binary polynomial p(2)
Implications for a 11 = aPC2)
p ~ 1//the leading digit is always 1 for n 2: 1
fori~ I
-1 downto 0 do
p+--2p+b;
fori~ l-1downto 0 do
aP ~ a2p+b;
r.~.'.~
i
!
!,
i'
234
Transform-and-Conquer
But
a2p+b; = a2p . ab; = (aP)2 . ab; =
I
2
(aP)
(aP) 2 . a
if b; = 0 .
if b; = 1
Thus, after initializing the accumulator's value to a, we can scan the bit string
representing the exponent to always square the last value of the accumulator and,
if the current binary digit is 1, also to multiply it by a. These observations lead to
the following left-to-right binary exponentiation method of computing a".
ALGORITHM
LeftRightBinaryExponentiation(a, b(n))
//Computes a" by the left-to-right binary exponentiation algorithm
//Input: A number a and a list b(n) of binary digits b 1 , ... , b0
II
in the binary expansion of a positive integer n
//Output: The value of a"
product +-- a
for i <-- I - 1 downto 0 do
product +--product* product
if b; = 1 product +--product * a
return product
EXAMPLE 2 Compute a 13 by the left-to-right binary exponentiation algorithm.
Here, n = 13 = 1101 2 . So we have
binary digits of n
1
product accumulator
a
Ill
Since the algorithm makes one or two multiplications on each repetition of its
only loop, the total number of multiplications M (n) made by it in computing a" is
(b- 1) <( M(n) <( 2(b- 1),
where b is the length of the bit string representing the exponent n. Taking into
account that b- 1 = Llog2 nj, we can conclude that the efficiency of the leftto-right binary exponentiation is logarithmic. Thus, this algorithm is in a better
efficiency class than the brute-force exponentiation, which always requires n - 1
multiplications.
The right-to-left binary exponentiation uses the same binary polynomial p(2)
(see (6.4)) yielding the value of n. But rather than applying Horner's rule to it as
the previous method did, this one exploits it differently:
6.5
Horner's Rule and Binary Exponentiation
235
Thus, a" can be computed as the product of the terms
if h; = 1 ,
Ifb; = 0
that is, the product of consecutive terms a 2;, skipping those for which the binary
digit h; is zero. In addition, we can compute az; by simply squaring tbe same term
we computed for the previous value of i since az; ~ (ai-'p. So we compute all such
powers of a from the smallest to the largest (from right to left), but we include in
the product accumulator only those whose corresponding binary digit is 1. Here
is a pseudocode of this algorithm.
AlGORITHM
RightLeftBinaryExponentiation(a, b(n))
//Computes a" by the right-to-left binary exponentiation algorithm
//Input: A number a and a list b(n) of binary digits b 1 , ... , b0
II
in the binary expansion of a nonnegative integer n
//Output: The value of a"
term +--a //initializes a 2;
if b0 = 1 product +--a
else product +-- 1
for i +-- 1 to I do
term+- term* term
if b; = 1 product +--product* term
return product
EXAMPLE 3 Compute a 13 by the right-to-left binary exponentiation method.
Here, n = 13 = 1101 2 . So we have the following table filled in from right to left.
1
0
a'
a'
a5·a8=al3
a
a
binary digits of n
terms ai
product accumulator
1111
Obviously, the algorithm's efficiency is also logarithmic for the same reason
the left-to-right binary multiplication is. The usefulness of both binary exponentiation algorithms is reduced somewhat by their reliance ou availability of the explicit
binary expansion of exponent n. Problem 8 in the exercises asks you to design an
algorithm that does not have this shortcoming.
236
Transform-and-Conquer
-~~~-Exercises
1. Consider the following brute-force algorithm for evaluating a polynomial.
ALGORITHM
BruteForcePolynomia!Evaluation(P[O.. n], x)
//The algorithm computes the value of polynomial P at a given point x
//by the "highest to lowest term" brute-force algorithm
//Input: An array P[O .. n] of the coefficients of a polynomial of degree n,
II
stored from the lowest to the highest and a number x
//Output: The value of the polynomial at the point x
p <-- 0.0
for i +-- n downto 0 do
power +-1
for j <-- 1 to i do
power+- power* x
p +-- p + P[i] *power
return p
Find the total number of multiplications and the total number of additions
made by this algorithm.
2. Write a pseudocode for the brute-force polynomial evaluation that stems from
substituting a given value of the variable into the polynomial's formula and
evaluating it from the lowest term to the highest one. Determine the number
of multiplications and the number of additions made by this algorithm.
3. a. Estimate how much faster Horner's rule is compared to the "lowest-tohighest term" brute-force algorithm of Problem 2 if (i) the time of one
multiplication is significantly larger than the time of one addition; (ii) the
time of one multiplication is about the same as the time of one addition.
b. Is Horner's rule more time efficient at the expense of being less space
efficient than the brute-force algorithm?
4. a. Apply Horner's rule to evaluate the polynomial
p(x) = 3x 4 - x 3 + 2x + 5 at x = -2.
b. Use the results of the above application of Horner's rule to fmd the quotient and remainder of the division of p(x) by x + 2.
5. Compare the number of multiplications and additions/subtractions needed
by the "long division" of a polynomial p(x) = a,x" + a,_ 1x"- 1 + · · · + a0 by
x - c, where c is some constant, with the number of these operations in
"synthetic division."
6. a. Apply the left-to-right binary exponentiation algorithm to compute a 17
6.6
Problem Reduction
237
b. Is it possible to extend the left-to-right binary exponentiation algorithm to
work for every nonnegative integer exponent?
7. Apply the right-to-left binary exponentiation algorithm to compute a 17
@ Design a nonrecursive algorithm for computing a" that mimics the right -to-left
binary exponentiation but does not explicitly use the binary representation of
n.
@Is it a good idea to use a general-purpose polynomial-evaluation algorithm
such as Horner's rule to evaluate the polynomial p(x) = x" + x"- 1 + · ·. +
X+ 1?
10. According to the corollary of the Fundamental Theorem of Algebra, every
·-· polynomial
p ( x ) =a11 x n + a11 _tX n-1 + · · · +ao
can be represented in the form
p(x) = a,(x- Xj)(x- x2) ... (x- x,)
where x1, x 2 , ... , x, are the roots of the polynomial (generally, complex and
not necessarily distinct). Discuss which of the two representations is more
convenient for each of the following operations:
a. polynomial evaluation at a given point
b. addition of two polynomials
c. multiplication of two polynomials
11. Polynomial puzzle Find the polynomial p(x) of degree eight such that p(i) =
1/ i fori = 1, 2, ... , 9.
6.6 Problem Reduction
Here is my version of a well-known joke about mathematicians. Professor X, a
noted mathematician, noticed that when his wife wanted to boil water for their
tea, she took their kettle from a kitchen cabinet, filled it with water, and put it
on the stove. Once, when his wife was away (if you have to know, she was signing
her best-seller in a local bookstore), the professor had to boil water by himself.
He saw that the kettle was sitting on the kitchen counter. What did Professor X
do? He put the kettle in the cupboard first and then proceeded to follow his wife's
routine.
The way Professor X approached his task is an example of an important
problem-solving strategy called problem reduction. If you need to solve a problem,
reduce it to another problem that you know how to solve (Figure 6.15).
238
Transform-and-Conquer
alg.A
reduction
Problem 1
(to be solved)
Problem 2
(solvable by alg. A)
solution
to Problem 2
FIGURE 6.15 Problem reduction strategy
The joke about the professor notwithstanding, the idea of problem reduction
plays a central role in theoretical computer science, where it is used to classify
problems according to their complexity. We touch on this classification in Chapter
11. But the strategy can be used for actual problem solving, too. The practical
difficulty in applying it lies, of course, in finding a problem to which the problem
at hand should be reduced. Moreover, if we want our efforts to be of practical
value, we need our reduction-based algorithm to be more efficient than solving
the original problem directly.
Note that we have already encountered this technique earlier in the book.
In Section 6.5, for example, we mentioned the so-called synthetic division done
by applying Horner's rule for polynomial evaluation. In Section 4.6, we used
the following fact from analytical geometry: if Pt = Cxt, Ytl, P2 = (x2, Y2l. and
p 3 = (x 3 , y 3 ) are three arbitrary points in the plane, then the determinant
Xj
Yt
x2
Y2
XJ
YJ
1
1 = XtJ'2 + X3Y1 + X2Y3 - X3Y2- X2Y1- XJY3
1
is positive if and only if the point p 3 is to the left of the directed line p;p', through
points p 1 and p 2 • In other words, we reduced a geometric question about the
relative locations of three points to a question about the sign of a determinant.
In fact, the entire idea of analytical geometry is based on reducing geometric
problems to algebraic ones. And the vast majority of geometric algorithms take
advantage of this historic insight by Rene Descartes (1596-1650). In this section,
we give a few more examples of algorithms based on the strategy of problem
reduction.
Computing the Least Common Multiple
Recall that the least common multiple of two positive integers m and n, denoted
lcm(m, n ), is defined as the smallest integer that is divisible by both m and n. For
example, lcm(24, 60) ~ 120, and lcm(ll, 5) ~ 55. The least common multiple is
one of the most important notions in elementary arithmetic and algebra; perhaps
you remember the following middle-school method for computing it. Given the
prime factorizations of m and n, !cm(m, n) can be computed as the product of all
the common prime factors of m and n times the product of m 's prime factors that
are not inn times n 's prime factors that are not in m. For example,
l
6.6
239
Problem Reduction
24=2·2·2-3
60=2·2·3·5
lcm(24, 60) = (2 · 2. 3) . 2. 5 = 120.
As a computational procedure, this algorithm has the same drawbacks as the
middle-school algorithm for computing the greatest common divisor discussed
in Section 1.1: it is inefficient and requires a list of consecutive primes.
A much more efficient algorithm for computing the least common multiple
can be designed by using problem reduction. After all, there is a very efficient
algorithm (Euclid's algorithm) for finding the greatest common divisor, which is a
product of all the common prime factors of m and n. Can we find a formula relating
lcm(m, n) and gcd(m, n)? It is not difficult to see that the product oflcm(m, n) and
gcd(m, n) includes every factor of m and n exactly once and hence is simply equal
to the product of m and n. This observation leads to the formula
m·n
lcm(m, n) = -.,.--gcd(m, n)
where gcd(m, n) can be computed very efficiently by Euclid's algorithm.
Counting Paths in a Graph
As our next example, we consider the problem of counting paths between two
vertices in a graph. It is not difficult to prove by mathematical induction that the
number of different paths of length k > 0 from the ith vertex to the jth vertex
of a graph (undirected or directed) equals the (i, j)th element of Ak where A is
the adjacency matrix of the graph. (Incidentally, the exponentiation algorithms
we discussed before for computing powers of numbers are applicable to matrices
as well.) Thus, the problem of counting a graph's paths can be solved with an
algorithm for computing an appropriate power of its adjacency matrix.
As a specific example, consider the graph of Figure 6.16. Its adjacency matrix
A and its square A 2 indicate the number of paths of length 1 and 2, respectively,
between the corresponding vertices of the graph. In particular, there are three
paths oflength 2 that start and end at vertex a: a - b- a, a - c- a, and a - d -a,
but only one path of length 2 from a to c: a - d -c.
a
b
c
a
d
b
c
d
Is: T tl T i]
0
A=
d
b
c
d
1
1
1
0
0
0
0
0
A2=
b
c
d
0
1
1
2
FIGURE 6.16 A graph, its adjacency matrix A, and its square A 2 The elements of A and
A2 indicate the number of paths of lengths 1 and 2, respectively.
240
Transform-and-Conquer
Reduction of Optimization Problems
Our next example deals with solving optimization problems. If a problem asks to
find a maximum of some function, it is said to be a maximization problem; if it
asks to find a function's minimum, it is called a minimization problem. Suppose
now that you have to find a minimum of some function f (x) and you know an
algorithm for maximizing the function. How can you take advantage of the latter?
The answer lies in the simple formula
min f(x) =-max[- f(x)].
In other words, to minimize a function, we can maximize its negative instead and,
to get a correct minimal value of the function itself, change the sign of the answer.
This property is illustrated for a function of one real variable in Figure 6.17.
Of course, the formula
max f(x) =-min[- f(x)]
is valid as well; it shows how a maximization problem can be reduced to an
equivalent minimization problem.
This relationship between minimization and maximization problems is very
general: it holds for functions defined on any domain D. In particular, we can
v
fix)
fix*) --------------~--,..-=-- - " '
x*
---+------------~--------------------------~X
-fix*) -------- -------c-.oc--~--'--
-fix)
FIGURE 6.17 Relationship between minimization and maximization problems:
min f(x) =-max[- f(x)]
6.6
Problem Reduction
241
apply it to functions of several variables subject to additional constraints. A
very important class of such problems is introduced in the next portion of this
section.
Now that we are on the topic of function optimization, it is worth pointing out
that the standard calculus procedure for finding extremum points of a function is,
in fact, also based on problem reduction. Indeed, it suggests finding the function's
derivative J'(x) and then solving the equation f'(x) = 0 to find the function's
critical points. In other words, the optimization problem is reduced to the problem
of solving an equation as the principal part of finding extremum points. Note
that we are not calling the calculus procedure an algorithm, since it is not clearly
defined. In fact, there is no general method for solving equations. A little secret of
calculus textbooks is that problems are carefully selected so that critical points
can always be found without difficulty. This makes the lives of both students
and instructors easier but, in the process, may unintentionally create a wrong
impression in students' minds.
linear Programming
Many problems of optimal decision making can be reduced to an instance of the
linear programming problem, which is a problem of optimizing a linear function
of several variables subject to constraints in the form of linear equations and linear
inequalities.
EXAMPLE 1 Consider a university endowment that needs to invest $100 million.
This sum must be split between three types of investments: stocks, bonds, and cash.
The endowment managers expect an annual return of 10%,7%, and 3% for their
stock, bond, and cash investments, respectively. Since stocks are more risky than
bonds, the endowment rules require the amount invested in stocks to be no more
than one third of the moneys invested in bonds. In addition, at least 25% of the
total amount invested in stocks and bonds must be invested in cash. How should
the managers invest the money to maximize the return?
Let us create a mathematical model of this problem. Let x, y, and z be the
amounts (in millions of dollars) invested in stocks, bonds, and cash, respectively.
By using these variables, we can pose the following optimization problem:
maximize 0.10x + 0.07y + 0.03z
subject to x + y + z = 100
1
-y
- 3
X <
z 2: 0.25(x
+ y)
X 2: 0, y 2: 0, Z 2: 0.
1111
242
Transform-and-Conquer
Although this particular problem is both small and simple, it does show how a
problem of optimal decision making can be reduced to an instance of the general
linear programming problem
maximize (or minimize)
subject to
+ · · · + C11 X 11
ailx 1 + · · · + ai X
CJXJ
11
11 ::::;
(or > or=) bi
fori= 1, ... , m
Xt :::. 0, ... , X 11 :::_ 0.
(The last group of constraints-called the nonnegativity constraints-are, strictly
speaking, unnecessary because they are special cases of more general constraints
ailxt + · · · + ai 11 x 11 ~hi, but it is convenient to treat them separately.)
Linear programming has proved to be flexible enough to model a wide variety
of important applications, such as airline crew scheduling, transportation and
communication network planning, oil exploration and refining, and industrial
production optimization. In fact, linear programming is considered by many as
one of the most important achievements in the history of applied mathematics.
The classic algorithm for this problem is called the simplex method (Section 10.1). It was discovered by the U.S. mathematician George Dantzig in the
1940s [Dan63]. Although the worst-case efficiency of this algorithm is known to
be exponential, it performs very well on typical inputs. The efforts of many computer scientists over the last 50 years have polished the algorithm and its computer
implementations to a point where problems with tens if not hundreds of thousands
of variables and constraints can be solved in a reasonable amount of time. More-
over, a few other algorithms for solving the general linear programming problem
have been discovered relatively recently; the best-known of them is by Narendra Karmarkar [Kar84]. The theoretical advantage of these newer algorithms lies
in their proven polynomial worst -case efficiency; Karmarkar's algorithm has also
been found to be competitive with the simplex method in empirical tests.
It is important to stress, however, that the simplex method and Karmarkar's
algorithm can successfully handle only linear programming problems that do not
limit its variables to integer values. When variables of a linear programming problem are required to be integers, the linear programming problem is said to be an
integer linear programming problem. Integer linear programming problems are
known to be much more difficult. There is no known polynomial-time algorithm
for solving an arbitrary instance of the general integer linear programming problem and, as we see in Chapter 11, such an algorithm quite possibly does not exist.
Other approaches such as the branch-and-bound technique discussed in Section
12.2 are typically used for solving integer linear programming problems.
EXAMPLE 2 Let us see how the knapsack problem can be reduced to a linear programming problem. Recall from Section 3.4 that the knapsack problem
can be posed as follows. Given a knapsack of capacity W and n items of weights
w 1, ... , wn and values v1, ... , V1l' find the most valuable subset of the items
that fits into the knapsack. We consider first the continuous (or fractional) ver-
6.6
Problem Reduction
243
sion of the problem, in which any fraction of any given item can be taken into
the knapsack. Let x 1, j = 1, ... , 11, be a variable representing a fraction of item
j taken into the knapsack. Obviously, x 1 must satisfy the inequality 0 S x s
1
1. Then the total weight of the selected items can be expressed by the sums
Lj=1 w.;x1, and their total value by the sum Lj=l v1xJ. Thus, the continuous ver~
sion of the knapsack problem can be posed as the following linear programming
problem:
maximize
subject to
"
"W·X· < W
.i.....J J J }=1
for j = 1, ... , 11.
There is no need to apply a general method for solving linear programming
problems here: this particular problem can be solved by a simple special algorithm
that is introduced in Section 12.3. (But why wait? Try to discover it on your
own now.) This reduction of the knapsack problem to an instance of the linear
programming problem is still useful, though, to prove the correctness of the
algorithm in question.
In the discrete (or 0-1) version of the knapsack problem, we are allowed either
to take an item in its entirety or not to take it at all. Hence, we have the following
integer linear programming problem for this version:
maximize
"
Lv.ixJ
}=1
"
subject to
"W·X· <
X; E {0, 1}
for j = 1, ... , n.
L...t.!J}=l
W
This seemingly minor modification makes a drastic difference for the complexity of this and similar problems constrained to take only discrete values in
their potential ranges. Despite the fact that the 0-1 version might seem to be easier because it can ignore any subset of the continuous version that has a fractional
value of an item, the 0-1 version is, in fact, much more complicated than its con-
tinuous counterpart. The reader interested in specific algorithms for solving this
problem will find a wealth of literature on the subject, including the monographs
[Mar90] and [Ke104].
Ill
__,..
I
244
Transform-and-Conquer
Reduction to Graph Problems
As we mentioned in Section 1.3, many problems can be solved by a reduction
to one of the standard graph problems. This is true in particular for a variety of
puzzles and games. In these applications, vertices of a graph typically represent
possible states of the problem in question while edges indicate permitted transitions among such states. One of the graph's vertices represents an initial state,
while another represents a goal state of the problem. (There might be several
vertices of the latter kind.) Such a graph is called a state-space graph. Thus, the
transformation just described reduces the problem to the question about a path
from the initial-state vertex to a goal-state vertex.
EXAMPLE As a specific example, let us revisit the classic river-crossing puzzle
of Problem 1 in the exercises to Section 1.2. A peasant finds himself on a river
bank with a wolf, a goat, and a head of cabbage. He needs to transport all three
to the other side of the river in his boat. However, the boat has room only for the
peasant himself and one other item (either the wolf, the goat, or the cabbage). In
his absence, the wolf would eat the goat, and the goat would eat the cabbage. Find
a way for the peasant to solve his problem or prove that it has no solution.
The state-space graph for this problem is given in Figure 6.18. Its vertices are
labeled to indicate the states they represent: P, w, g, c stand for the peasant, the
wolf, the goat, and the cabbage, respectively; the two bars II denote the river;
Pg
Pwgc I I )--~'------( we I I Pg
p
c II Pwg
p
Pg
Pg I I we 1-----'-'"-----( II Pwge
FIGURE 6.18 State-space graph for the peasant. wolf, goat, and cabbage puzzle
6.6
Problem Reduction
245
for convenience, we also label the edges by indicating the boat's occupants for
each crossing. In terms of this graph, we are interested in frnding a path from the
initial-state vertex, labeled Pwgc II, to the final-state vertex, labeled II Pwgc.
It is easy to see that there exist two distinct simple paths from the initialstate vertex to the final state vertex (what are they?). If we find them by applying
breadth-first search, we get a formal proof that these paths have the smallest
number of edges possible. Hence, this puzzle has two solutions, each of which
requires seven river crossings.
Ill
Our success in solving this simple puzzle should not lead you to believe that
generating and investigating state-space graphs is always a straightforward task.
To get a better appreciation of them, consult books on artificial intelligence (AI),
the branch of computer science in which state-space graphs are a principal subject.
In this book, we deal with an important special case of state-space graphs in
Sections 12.1 and 12.2.
~~---Exercises 6.6 ~-------------1. a. Prove the equality
m ·n
lcm(m, n) = - - - gcd(m, 11)
that underlies the algorithm for computing lcm(m, n).
b. Euclid's algorithm is known to be in O(log 11). If it is the algorithm that is
used for computing gcd(m, 11), what is the efficiency of the algorithm for
computing lcm(m, 11)?
2. You are given a list of numbers for which you need to construct a min-heap.
(A min-heap is a complete binary tree in which every key is less than or equal
to the keys in its children.) How would you use an algorithm for constructing
a max-heap (a heap as defined in Section 6.4) to construct a min-heap?
3. Prove that the number of different paths oflength k > 0 from the ith vertex to
the jth vertex in a graph (undirected or directed) equals the (i, j)th element
of A k where A is the adjacency matrix of the graph.
4. a. Design an algorithm with a time efficiency better than cubic for checking
whether a graph with 11 vertices contains a cycle of length 3. [Man89],
p.326
b. Consider the following algorithm for the same problem. Starting at an arbitrary vertex, traverse the graph by deptb-first search and check whether its
depth-first search forest has a vertex with a back edge leading to its grandparent. If it does, the graph contains a triangle; if it does not, the graph
does not contain a triangle as its subgraph. Is this algorithm correct?
,'!'
il
246
Transform-and-Conquer
5. Given n > 3 points P1 = (x 1, Yl), ... , P, = (x,., y,) in the coordinate plane,
design an algorithm to check whether all the points lie within a triangle with
its vertices at three of the points given. (You can either design an algorithm
from scratch or reduce the problem to another one with a known algorithm.)
6. Consider the problem of finding, for a given positive integer n, the pair of
integers whose sum is n and whose product is as large as possible. Design an
efficient algorithm for this problem and indicate its efficiency class.
7. The assignment problem, introduced in Section 3.4, can be stated as follows.
There are n people who need to be assigned to execute n jobs, one person
per job. (That is, each person is assigned to exactly one job and each job is
assigned to exactly one person.) The cost that would accrue if the ith person is
assigned to the jthjob is a known quantity C[i, j] for each pair i, j = 1, ... , n.
The problem is to assign the people to the jobs to minimize the total cost of
the assignment. Express the assignment problem as a O-J linear programming
problem.
8. Solve the instance of the linear programming problem given in Section 6.6
maximize
0.1 Ox + 0.07y + 0.03z
subject to x + y + z = 100
1
X< -y
- 3
z 2: 0.25(x + y)
X 2: 0,
y 2: 0, Z 2: 0.
9. The graph-coloring problem is usually stated as the vertex-coloring problem:
assign the smallest number of colors to vertices of a given graph so that no
two adjacent vertices are the same color. Consider the edge-coloring problem:
assign the smallest number of colors possible to edges of a given graph so that
no two edges with the same endpoint are the same color. Explain how the
edge-coloring problem can be reduced to a vertex-coloring problem.
10. Consider the two-dimensional post office location problem: given n points
(xlo y 1) ... , (x,, y,) in the Cartesian plane, find a location (x, y) for a post
office that minimizes* I:;'~ 1 (lx;- xl + IY;- yl), the average Manhattan distance from the post office to the given points. Explain how this problem can
be efficiently solved by the problem-reduction technique, provided the post
office does not have to be located at one of the input points.
11. Jealous husbands There are n (n 2: 2) married couples who need to cross a
river. They have a boat that can hold no more than two people at a time. To
complicate matters, all the husbands are jealous and will not agree on any
crossing procedure that would put a wife on the same bank of the river with
another woman's husband without the wife's husband being there too, even if
Summary
247
there are other people on the same bank. Can they cross the river under such
constraints?
a. Solve the problem for n = 2.
b. Solve the problem for n = 3, which is the classical version of this problem.
c. Does the problem have a solution for every n 2: 4? If it does, indicate how
many river crossings it will take; if it does not, explain why.
SUMMARY
"
Transform-and-conquer is the fourth general algorithm design (and problemsolving) strategy discussed in the book. It is, in fact, a group of techniques
based on the idea of transformation to a problem that is easier to solve.
"
There are three principal varieties of the transform-and-conquer strategy:
instance simplification, representation change, and problem reduction.
"
Instance simplification is a technique of transforming an instance of a problem
to an instance of the same problem with some special property that makes the
problem easier to solve. List presorting, Gaussian elimination, and rotations
in AVL trees are good examples of this technique.
"
Representation change implies changing one representation of a problem's
instance into another representation of the same instance. Examples discussed
in this chapterinclude representation of a set by a 2-3 tree, heaps and heapsort,
Horner's rule for polynomial evaluation, and two binary exponentiation
algorithms.
"'
Problem reduction calls for transforming a given problem to another problem
that can be solved by a known algorithm. Among examples of applying this
idea to algorithmic problem solving (see Section 6.6), reductions to linear
programming and reductions to graph problems are especially important.
"
Some examples used to illustrate the transform-and-conquer techniques
happen to be very important data structures and algorithms. They are heaps
and heapsort, AVL and 2-3 trees, Gaussian elimination, and Horner's rule.
"
A heap is an essentially complete binary tree with keys (one per node)
satisfying the parental dominance requirement. Though defined as binary
trees, heaps are normally implemented as arrays. Heaps are most important
for the efficient implementation of priority queues; they also underlie
heapsort.
"
Heapsort is a theoretically important sorting algorithm based on arranging
elements of an array in a heap and then successively removing the largest
~.· ~
~- ' i
r.·
. .i. ....,;.
..
248
Transform-and-Conquer
element from a remaining heap. The algorithm's running time is in El(n log n)
both in the worst case and in the average case; in addition, it is in place.
m
AVL trees are binary search trees that are always balanced to the extent
possible for a binary tree. The balance is maintained by transformations of
four types called rotations. All basic operations on AVL trees are in O(log n);
it eliminates the bad worst-case efficiency of classic binary search trees.
m
2-3 trees achieve a perfect balance in a search tree by allowing a node to
contain up to two ordered keys and have up to three children. This idea can
be generalized to yield very important B-trees, discussed later in the book.
WJ
Gaussian elimination-an algorithm for solving systems of linear equationsis a principal algorithm in linear algebra. It solves a system by transforming it
to an equivalent system with an upper-triangular coefficient matrix, which is
easy to solve by backward substitutions. Gaussian elimination requires about
~n 3 multiplications.
"
Horner's rule is an optimal algorithm for polynomial evaluation without
coefficient preprocessing. It requires only 11 multiplications and 11 additions
to evaluate an 11-degree polynomial at a given point. Horner's rule also has a
few useful by-products such as the synthetic division algorithm.
6
Two binary exponentiation algorithms for computing a 11 were introduced in
Section 6.5. Both of them exploit the binary representation of the exponent
n, but they process it in the opposite directions: left to right and right to left.
11
. j
Linear programming concerns optimizing a linear function of several variables subject to constraints in the form oflinear equations and linear inequalities. There are efficient algorithms capable of solving very large instances
of this problem with many thousands of variables and constraints, provided
the variables are not required to be integers. The latter, called integer linear
programming, constitute a much more difficult class of problems .
Space and Time Tradeoffs
Things which matter most must never be at the mercy of things which matter
less.
-Johann Wolfgang von Goethe (1749-1832)
S
pace and time tradeoffs in algorithm design are a well-known issue for both
theoreticians and practitioners of computing. Consider, as an example, the
problem of computing values of a function at many points in its domain. If it is
time that is at a premium, we can precompute the function's values and store them
in a table. This is exactly what human computers had to do before the advent of
electronic computers, in the process burdening libraries with thick volumes of
mathematical tables. Though such tables have lost much of their appeal with the
widespread use of electronic computers, the underlying idea has proven to be quite
useful in the development of several important algorithms for other problems.
In somewhat more general terms, the idea is to preprocess the problem's input,
in whole or in part, and store the additional information obtained to accelerate
solving the problem afterward. We call this approach input enhancement 1 and
discuss the following algorithms based on it:
Ill
"
counting methods for sorting (Section 7.1)
Boyer-Moore algorithm for string matching and its simplified version suggested by Horspool (Section 7.2)
The other type of technique that exploits space-for-time tradeoffs simply uses
extra space to facilitate faster and/or more flexible access to the data. We call this
approach prestructuring. This name highlights two facets of this variation of the
space-for-time tradeoff: some processing is done before a problem in question
---·---1.
TilC standard terms used synonymously for this technique arc preprocessing and preconditioning.
Confusingly, these terms can also be applied to methods that use the idea of preprocessing but do not
use extra space (see Chapter 6). Thus, in order to avoid confusion, we use "input enhancement" as a
special name for the space-for-time tradeoff technique being discussed here.
249
-.,.
I
250
Space and 1ime Tradeoffs
is actually solved but, unlike the input-enhancement variety, it deals with access
structuring. We illustrate this approach by
,.
,.
hashing (Section 7.3)
indexing with B-trees (Section 7.4)
There is one more algorithm design technique related to the space-for-time
tradeoff idea: dynamic pl'ogl'amming. Tbis strategy is based on recording solutions to overlapping subproblems of a given problem in a table from which a solution to the problem in question is then obtained. We discuss this well-developed
technique separately, in the next chapter of the book.
Two final comments about the interplay between time and space in algorithm
design need to be made. First, the two resources-time and space-do not have
to compete with each other in all design situations. In fact, they can align to
bring an algorithmic solution that minimizes both the running time and the space
consumed. Such a situation arises, in particular, when an algorithm uses a spaceefficient data structure to represent a problem's input, which leads, in turn, to
a faster algorithm. Consider, as an example, the problem of traversing graphs.
Recall that the time efficiency of the two principal traversal algorithms-depthfirst search and breadth-first search-depends on the data structure used for
representing graphs: it is El(n 2 ) for the adjacency matrix representation and EJ(n +
m) for the adjacency list representation, where n and m are the numbers of
vertices and edges, respectively. If input graphs are sparse, i.e., have few edges
relative to the number of vertices (say, m E 0 (n) ), the adjacency list representation
may well be more efficient from both the space and the running-time points of
view. The same situation arises in the manipulation of sparse matrices and sparse
polynomials: if the percentage of zeros in such objects is sufficiently high, we can
save both space and time by ignoring zeros in the objects' representation and
processing.
Second, one cannot discuss space-time tradeoffs without mentioning the
hugely important area of data compression. Note, however, that in data compression, size reduction is the goal rather than a technique for solving another
problem. We discuss just one data compression algorithm, in the next chapter.
The reader interested in this topic will find a wealth of algorithms in such books
as [SayOO].
1 Sorting by Counting
As a first example of applying the input enhancement technique, we discuss its
application to the sorting problem. One rather obvious idea is to count, for each
element of a list to be sorted, the total number of elements smaller than this
element and record the results in a table. These numbers will indicate the positions
of the elements in the sorted list: e.g., if the count is 10 for some element, it should
be in the 11th position (with index 10, if we start counting with 0) in the sorted
7.1
Sorting by Counting
251
array. Thus, we will be able to sort the list by simply copying its elements to their
appropriate positions in a new, sorted list. This algorithm is called comparison
counting sort (Figure 7.1).
AlGORITHM
ComparisonCountingSort(A[O .. n - 1])
//Sorts an array by comparison counting
//Input: An array A[O .. n - 1] of orderable elements
//Output: Array S[O .. n - 1] of A's elements sorted in non decreasing order
fori <-- 0 ton - 1 do Count[i] <-- 0
for i <-- 0 to n - 2 do
for j <-- i + 1 to n - 1 do
if A(i] < A[Jl
Count[}] <-- Count[}]+ 1
else Count[i] <-- Count[i] + 1
fori +- 0 ton - 1 do S[ Count[i]] <-- A[i]
returnS
What is the time efficiency of this algorithm? It should be quadratic because
the algorithm considers all the different pairs of ann-element array. More formally,
the number of times its basic operation, the comparison A[i] < A[j], is executed
is equal to the sum we have encountered several times already:
n-2 n-1
n-2
n-2
"""""' L....
"""""' 1= """""'
.
n(n- 1) .
C(n)= L....
L....[(n -1) -(r. +1)+1]= """""'
L._.(n -1-r)=
i=O }=i+l
i=O
i=O
2
Since the algorithm makes the same number of key_t;_~_s as selection sort
and in addition uses a linear amount of extra space, it can hardly be rec~d
for practical use.
Array A10 .. 5)
Initially
After pass i = 0
After pass i = 1
After pass i = 2
After pass i = 3
After pass i = 4
Final state
Array 5[0 .. 51
FIGURE 7.1
I 62 I 31 I s4 1 96 1 19 1 47 1
Count[)
Count II
Count[)
Count II
Count[)
Count II
Count[)
0
3
0
0
1
0
1
2
4
3
1
4
0
1
2
3
5
5
0
0
0
0
0
0
0
0
0
1
1
1
2
2
1 19 1 31 1 47 1 62 1 s4 1 96 I
Example of sorting by comparison counting
252
Space and lime Tradeoffs
But the couuting idea does work productively in a situation in which elements
to be sorted belong to a known small set of values. Assume, for example, that
we have to sort a list whose values can be either 1 or 2. Rather than applying a
general sorting algorithm, we should be able to take advantage of this additional
information about values to be sorted. Indeed, we can scan the list to compute
the number of 1's and the number of 2's in it and then, on the second pass,
simply make the appropriate number of the first elements equal to 1 and the
remaining elements equal to 2. More generally, if element values are integers
between some lower bound I and upper bound u, we can compute the frequency
of each of those values and store them in array F[O .. u - 1]. Then the first F[O]
positions in the sorted list must be filled with I, the next F[1] positions with I + 1,
and so on. All this can be done, of course, only if we can overwrite the given
elements.
Let us consider a more realistic situation of sorting a list of items with some
other information associated with their keys so that we cannot overwrite the list's
elements. Then we can copy elements into a new array S[O .. n - 1] to hold the sorted
list as follows. The elements of A whose values are equal to the lowest possible
value/ are copied into the first F[O] elements of S, i.e., positions Othrough F[O]- 1,
the elements of value I+ 1 are copied to positions from F[O] to (F[O] + F[1]) - 1,
and so on. Since such accumulated sums of frequencies are called a distribution
in statistics, the method itself is known as distribution counting.
EXAMPLE
Consider sorting the array
I 13 111 112 113 I 12
12
whose values are known to come from the set {11, 12. 13} and should not be
overwritten in the process of sorting. The frequency and distribution arrays are
as follows:
Array values
11
12
l3
Frequencies
Distribution values
1
1
3
4
2
6
Note that the distribution values indicate the proper positions for the last occurrences of their elements in the final sorted array. If we index array positions from 0
ton - 1, the distribution values must be reduced by 1 to get corresponding element
positions.
It is more convenient to process the input array right to left. For the example,
the last element is 12, and, since its distribution value is 4, we place this 12 in
position 4 - 1 = 3 of the array S that will hold the sorted list. Then we decrease
7.1
A 151 = 12
1
0[0.2]
4
6
Al41=12
1
3
A [3] = 13
A [21 = 12
A 111 = 11
A [O[ = 13
1
1
2
1
0
2
1
1
6
6
5
5
5
Sorting by Counting
253
5[0 5]
12
12
13
12
11
13
FIGURE 7.2 Example of sorting by distribution counting. The distribution values being
decremented are shown in bold.
the 12's distribution value by 1 and proceed to the next (from the right) element
in the given array. The entire processing of this example is depicted in Figure 7.2.
Ill
Here is a pseudocode of this algorithm.
ALGORITHM
DistributionCounting(A[O.. n- 1], I, u)
//Sorts an array of integers from a limited range by distribution counting
//Input: An array A[O .. n - 1] of integers between l and u (I :" u)
//Output: Array S[O .. n - 1] of A's elements sorted in nondecreasing order
//initialize frequencies
for j +- 0 to u -I do D[j] +- 0
fori +- 0 ton - 1 do D[ A[i]- l] +- D[ A[i]- 1] + 1 //compute frequencies
//reuse for distribution
for j +- 1 to u - l do D[j] +- D[j - 1] + D[Jl
fori+- n -1 downto 0 do
j +- A[i]-1
S[ D[j]- 1] +- A[i]
D[j] +- D[j]- 1
returnS
Assuming that the range of array values is fixed, this is obviously a linear
algorithm because it makes just two consecutive passes through its input array
A. This is a better time-efficiency class than that of the most efficient sorting
algorithms-mergesort, quicksort, and heapsort-we have encountered. It is important to remember, however, that this efficiency is obtained by exploiting the
specific nature of input lists on which sorting by distribution counting works, in
addition to trading space for time.
-----Exercises 7.1 - - - - - - - - - - - - - - - 1. Is it possible to exchange numeric values of two variables, say, u and v, without
using any extra storage?
254
Space and Time Tradeoffs
2. Will the comparison counting algorithm work correctly for arrays with equal
values?
3. Assuming that the set of possible list values is {a, b, c, d}, sort the following
list in alphabetical order by the distribution counting algorithm:
b, c, d, c, b, a, a, b.
4. Is the distribution counting algorithm stable?
5. Design a one-line algorithm for sorting any array of size n whose values are n
distinct integers from 1 ton.
6. The ancestry problem asks to determine whether a vertex u is an ancestor
of vertex v in a given binary (or, more generally, rooted ordered) tree of n
vertice& Design a O(n) input enhancement algorithm that provides sufficient
information to solve this problem for any pair of the tree's vertices in constant
time.
7. The following technique, known as virtual initialization, provides a timeefficient way to initialize just some elements of a given array A[O .. n - 1] so
that for each of its elements, we can say in constant time whether it has been
initialized and, if it has been, with which value. This is done by utilizing a
variable counter for the number of initialized elements in A and two auxiliary
arrays of the same size, say B[O .. n -1] and C[O .. n -1], defined as follows.
B[O], ... , B[counter -1] contain the indices of the elements of A that were
initialized: B [0] contains the index of the element initialized first, B [1] contains
the index of the element initialized second, and so on. Furthermore, if A[i]was
the kth element (0 :S k :S counter - 1) to be initialized, C[i] contains k.
a. Sketch the state of arrays A[0 .. 7], B[0 .. 7], and C[0 .. 7] after the three assignments
A[3] +- x; A[7] +- z; A[l] +- y.
b. In general, how can we check with this scheme whether A[i] has been
initialized and, if it has been, with which value?
8. a. Write a program for multiplying two sparse matrices, a p-by-q matrix A
and a q-by-r matrix B.
b. Write a program for multiplying two sparse polynomials p(x) and q(x) of
degrees m and n, respectively.
9. Tic-tac-toe Write a program that plays the game of tic-tac-toe with the human user by storing all possible positions on the game's 3-by-3 board along
with the best move for each of them.
7.2
7
Input Enhancement in String Matching
255
Input Enhancement in String Matching
In this section, we see how the technique of input enhancement can be applied
to the problem of string matching. Recall that the problem of string matching
requires finding an occurrence of a given string of m characters (called the pattern)
in a longer string of n characters (called the text). We discussed the brute-force
algorithm for this problem in Section 3.2: it simply matches corresponding pairs
of characters in the pattern and the text left to right and, if a mismatch occurs,
shifts the pattern one position to the right for the next trial. Since the maximum
number of such trials is.Jz - m + l _ ~nd~ __}n the worst case, m comparisons need
to be made on each of them, the ·_worst-case__·number of character comparisons is
m(n -m + 1). This puts the worst-case performance of the brute-force algorithm
in the C"!(nm) class. On averag<,>, however, we should expect just a few comparisons
before a pattern\s shift, and for random natural-language texts, the average-case
efficiency indeed turns out to be in (11).
The good average-case efficiency of the brute-force algorithm can be considered both good and bad news. It is good news from a practical standpoint because
it makes the brute-force solution a passable candidate for practical applications
(especially for short patterns). It is bad news for a theoretician who would like to
invent a faster algorithm. Nevertheless, several better algorithms have been discovered. Most of them exploit the input enhancement idea: preprocess the pattern
to get some information about it, store this information in a table, and then use
this information during an actual search for the pattern in a given text. This is
exactly the idea behind the two best-known algorithms of this type: the KnuthMarris-Pratt algorithm [Knu77] and the Boyer-Moore algorithm [Boy77].
The principal difference between these two algorithms lies in the way they
compare characters of a pattern with their counterparts in a text: the KnuthMarris-Pratt algorithm does it left to right, while the Boyer-Moore algorithm does
it right to left. Since the latter idea leads to simpler algorithms, it is the only one
we will pursue here. (Note that the Boyer-Moore algorithm starts by aligning the
pattern against the beginning characters of the text; if the first trial fails, it shifts
the pattern to the right. It is comparisons within a trial tbat the algorithm does
right to left, starting with the last one in the pattern.)
Although the underlying idea of the Boyer-Moore algorithm is simple, its
actual implementation in a working method is less so. Therefore, we start our
discussion with a simplified version of the Boyer-Moore algorithm suggested by
R. Horspool [Hor80]. In addition to being simpler, Horspool's algorithm is not
necessarily less efficient than the Boyer-Moore algorithm on random strings.
"J
Horspool's Algorithm
Consider, as an example, searching for the pattern BARBER in some text:
c
BARBER
rr
r
!i
256
Space and lime Tradeoffs
Starting with the last R of the pattern and moving right to left, we compare the
corresponding pairs of characters in the pattern and the text. If all the pattern's
characters match successfully, a matching substring is found. (Then the search can
be either stopped altogether or continued if another occurrence of the same pattern is desired.) If, however, we encounter a mismatch, we need to shift the pattern
to the right. Clearly, we would like to make as large a shift as possible without risking the possibility of missing a matching substring in the text. Horspool's algorithm
determines the size of such a shift by looking at the character c of the text that
was aligned against the last character of the pattern. In general, the following four
possibilities can occur.
Case 1 If there are no c's in the pattern-e.g., c is letter S in our examplewe can safely shift the pattern by its entire length (if we shift less, some character
of the pattern would be aligned against the text's character c that is known not to
be in the pattern):
s
Jf
BARBER
BARBE R
Case 2 If there are occurrences of character c in the pattern but it is not the last
one there-e.g., c is letter B in our example-the shift should align the rightmost
occurrence of c in the pattern with the c in the text:
B
Jf
BARBE R
BARBER
Case 3 If c happens to be the last character in the pattern but there are no c's
among its other m - 1 characters, the shift should be similar to that of Case 1: the
pattern should be shifted by the entire pattern's length m, e.g.,
M E R
Jf II II
LEADER
LEADER
Case 4 Finally, if c happens to be the last character in the pattern and there
are other c's among its first m - 1 characters, the shift should be similar to that
of Case 2: the rightmost occurrence of c among the first m - 1 characters in the
pattern should be aligned with the text's c, e.g.,
7.2
Input Enhancement in String Matching
257
0 R
Jf II
REORDER
REORDER
These examples clearly demonstrate that right-to-left character comparisons
can lead to farther shifts of the pattern than the shifts by only one position
always made by the brute-force algorithm. However, if such an algorithm had
to check all the characters of the pattern on every trial, it would lose much
of this superiority. Fortunately, the idea of input enhancement makes repetitive
comparisons unnecessary. We can precompute shift sizes and store them in a table.
The table will be indexed by all possible characters that can be encountered in a
text, including, for natural language texts, the space, punctuation symbols, and
other special characters. (Note that no other information about the text in which
eventual searching will be done is required.) The table's entries will indicate the
shift sizes computed by the formula
the pattern's length m,
if c is not among the first m - 1 characters of the pattern
t(c) =
I
(7.1)
the distance from the rightmost c among the first m - 1 characters
of the pattern to its last character, otherwise
For example, for the pattern BARBER, all the table's entries will be equal to 6, except
for the entries forE, B, R, and A, which will be 1, 2, 3, and 4, respectively.
Here is a simple algorithm for computing the shift table entries. Initialize all
the entries to the pattern's length m and scan the pattern left to right repeating the
following step m - 1 times: for the jth character of the pattern (0 :S j :S m - 2),
overwrite its entry in the table with m - 1 - j, which is the character's distance to
the right end of the pattern. Note that since the algorithm scans the pattern from
left to right, the last overwrite will happen for a character's rightmost occurrenceexactly as we would like it to be.
ALGORITHM
ShiftTable(P[O .. m - 1])
//Fills the shift table used by Horspool's and Boyer-Moore algorithms
//Input: Pattern P[O .. m -1] and an alphabet of possible characters
//Output: Table[O .. size - 1] indexed by the alphabet's characters and
II
filled with shift sizes computed by formula (7.1)
initialize all the elements of Table with m
for j <-- 0 tom- 2 do Table[ P[j]] <-- m -1- j
return Table
258
Space and Time Tradeoffs
Now, we can summarize the algorithm as follows.
Horspool's algorithm
Step 1 For a given pattern of length m and the alphabet used in both the
pattern and text, construct the shift table as described above.
Step 2 Align the pattern against the beginning of the text.
Step 3 Repeat the following until either a matching substring is found or the
pattern reaches beyond the last character of the text. Starting with the
last character in the pattern, compare the corresponding characters in
the pattern and text until either all m characters are matched (then
stop) or a mismatching pair is encountered. In the latter case, retrieve
the entry t(c) from the c's column of the shift table where cis the text's
character currently aligned against the last character of the pattern,
and shift the pattern by t(c) characters to the right along the text.
Here is a pseudocode of Horspool's algorithm.
ALGORITHM
HorspoolMatching(P[O .. m -1], T[O .. n -1])
//Implements Horspool's algorithm for string matching
//Input: Pattern P[O .. m -1] and text T[O .. n -1]
//Output: The index of the left end of the first matching substring
II
or -1 if there are no matches
ShiftTable(P[O .. m - 1]) //generate Table of shifts
i ~m - 1
//position of the pattern's right end
while i :5: n - 1 do
k~0
//number of matched characters
while k s m- 1 and P[m- 1- k] = T[i- k] do
k ~k+1
ifk =m
retnrni-m+l
else i ~ i+Table[T[i]]
return -1
EXAMPLE As an example of a complete application of Horspool's algorithm,
consider searching for the pattern BARBER in a text that comprises English letters
and spaces (denoted by underscores). The shift table, as we mentioned, is filled as
follows:
character c
A
B
c
D
E
F
...
R
. ..
z
-
shift t(c)
4
2
6
6
1
6
6
3
6
6
6
7.2
Input Enhancement in String Matching
259
The actual search in a particular text proceeds as follows:
J I M
s A w ME I N A 8 AR8 ERS H0 P
8 A R 8 E R
8 A R 8 E R
8 A R 8 E R
8 A R 8 E R
8 A R 8 E R
8 A R 8 E R
1111
A simple example can demonstrate that the worst-case efficiency of Herspool's algorithm is in G(nm) (Problem 4 in the exercises). But for random texts,
it is in G(n), and, though in the same efficiency class, Horspool's algorithm is obviously faster on average than the brute-force algorithm. In fact, as mentioned,
it is often at least as efficient as its more sophisticated predecessor discovered by
R. Boyer and J. Moore.
Boyer-Moore Algorithm
Now we outline the Boyer-Moore algorithm itself. If the first comparison of the
rightmost character in the pattern with the corresponding character c in the text
fails, the algorithm does exactly the same thing as Horspool's algorithm. Namely,
it shifts the pattern to the right by the number of characters retrieved from the
table precomputed as explained earlier.
The two algorithms act differently, however, after some positive number k
(0 < k < m) of the pattern's characters are matched successfully before a mismatch
is encountered:
c
s 11 _1
Jf
Po
Pm-k-1
Pm-k
text
pattern
In this situation, the Boyer-Moore algorithm determines the shift size by considering two quantities. The first one is guided by the text's character c that caused
a mismatch with its counterpart in the pattern. Accordingly, it is called the badsymbol shift. The reasoning behind this shift is the reasoning we used in Barspool's algorithm. If cis not in the pattern, we shift the pattern to just pass this
c in the text. Conveniently, the size of this shift can be computed by the formula
t 1(c)- k, where t 1 (c) is the entry in the precomputed table used by Horspool's
algorithm and k is the number of matched characters:
so
Po
c
si-k+1
Jf
II
Pm-k-1
Pm-k
Po
s,
Sn-1
text
pattern
Pm-1
Pm-1
..
'Ff1
. . I,
260
Space and Time Tradeoffs
For example, if we search for the pattern BARBER in some text and match the
last two characters before failing on letter 5 in the text, we can shift the pattern
by It (5) - 2 = 6- 2 = 4 positions:
5 E R
Jf II II
BARBER
BARBER
The same formula can also be used when the mismatching character c of the
text occurs in the pattern, provided lt(c)- k > 0. For example, if we search for
the pattern BARBER in some text and match the last two characters before failing on letter A, we can shift the pattern by 11(A) - 2 = 4- 2 = 2 positions:
A E R
Jf II II
BARBER
BARBER
If 11(c) - k <( 0, we obviously do notwantto shift the pattern by 0 or a negative
number of positions. Rather, we can fall back on the brute-force thinking and
simply shift the pattern by one position to the right.
To summarize, the bad-symbol shift d 1 is computed by the Boyer-Moore
algorithm either as It (c) - k if this quantity is positive or as 1 if it is negative or
zero. This can be expressed by the following compact formula:
dt = max{lt(c)- k, 1].
(7.2)
The second type of shift is guided by a successful match of the last k > 0
characters of the pattern. We refer to the ending portion of the pattern as its suffix
of size k and denote it suff(k ). Accordingly, we call this type of shift the good-suffix
shift. We now apply the reasoning that guided us in filling the bad-symbol shift
table, which was based on a single alphabet character c, to the pattern's suffixes
of sizes 1, ... , m - 1 to fill in the good-suffix shift table.
Let us first consider the case when there is another occurrence of suff(k) in
the pattern or, to be more accurate, there is another occurrence of suff(k) not
preceded by the same character as in its last occurrence. (It would be useless to
shift the pattern to match another occurrence of suff(k) preceded by the same
character because this would simply repeat a failed trial.) In this case, we can
shift the pattern by the distance d2 between such a second rightmost occurrence
(not preceded by the same character as in the last occurrence) of suff(k) and its
rightmost occurrence. For example, for the pattern ABCBAB, these distances for
k = 1 and 2 will be 2 and 4, respectively:
7.2
Input Enhancement in String Matching
k
pattern
2
ABCBA!l_
ABCBAB
261
2
4
What is to be done if there is no other occurrence of sujf(k) not preceded by
the same character as in its last occurrence? In most cases, we can shift the pattern
by its entire length m. For example, for the pattern DBCBAB and k = 3, we can shift
the pattern by its entire length of 6 characters:
c B A B
lf II I II
DBCBAB
D B C B A B
Unfortunately, shifting the pattern by its entire length when there is no other
occurrence of suff(k) not preceded by the same character as in its last occurrence
is not always correct. For example, for the pattern ABCBAB and k = 3, shifting hy 6
could miss a matching substring that starts with the text's AB aligned with the last
two characters of the pattern:
cBABCBAB
lf II II II
ABCBAB
A B C B A B
Note that the shift by 6 is correct for the pattern DBCBAB but not for ABCBAB,
because the latter pattern has the same substring AB as its prefix (beginning part
of the pattern) and as its suffix (ending part of the pattern). To avoid such an
erroneous shift based on a suffix of size k, for which there is no other occurrence
in the pattern not preceded by the same character as in its last occurrence, we need
to find the longest prefix of size l < k that matches the suffix of the same size I.
If such a prefix exists, the shift size d 2 is computed as the distance between this
prefix and the corresponding suffix; otherwise, d2 is set to the pattern's length m.
As an example, here is the complete list of the d2 values-the good-suffix table of
the Boyer-Moore algorithm-for the pattern ABCBAB:
k
pattern
d2
1
2
3
ABCBA!l_
ABCBAB
ABCBAB
ABCBAB
ABCBAB
2
4
5
4
4
4
4
p
~-· .· · ·.·
' '
···r
..·'.·
.•·· ·
•..
•
;·:
I
...
262
Space and Time Tradeoffs
Now we are prepared to summarize the Boyer-Moore algorithm in its entirety.
The Boyer-Moore algorithm
Step 1 For a given pattern and the alphabet used in both the pattern and the
text, construct the bad-symbol shift table as described earlier.
Step 2 Using the pattern, construct the good-suffix shift table as described.
Step 3 Align the pattern against the beginning of the text.
Step 4 Repeat the following step until either a matching substring is found or
the pattern reaches beyond the last character of the text. Starting with
the last character in the pattern, compare the corresponding characters
in the pattern and the text until either all m character pairs are matched
(then stop) or a mismatching pair is encountered after k 2: 0 character
pairs are matched successfully. In the latter case, retrieve the entry
t1(c) from the c's column of the bad-symbol table where cis the text's
mismatched character. If k > 0, also retrieve the corresponding d 2
entry from the good-suffix table. Shift the pattern to the right by the
number of positions computed by the formula
ifk=O
if k > 0
(7.3)
where dt = max{t1(c)- k, 1].
Shifting by the maximum of the two available shifts when k > 0 is quite logical. The two shifts are based on the observations-the first one about a text's
mismatched character and the second one about a matched group of the pattern's
rightmost characters-that imply that shifting by less than d1 and d 2 characters, respectively, cannot lead to aligning the pattern with a matching substring in the text.
Since we are interested in shifting the pattern as far as possible without missing a
possible matching substring, we take the maximum of these two numbers.
EXAMPLE As a complete example, let us search for BAOBAB in a text made of
English letters and spaces. The bad-symbol table looks as follows:
6
······ytr··
I
7.2
B E S
S
B
B
A
0
E
w
A
B A
0
B A
K N
A
dt = lt(K)- 0 = 6
Input Enhancement in String Matching
B 0
263
u
T
B A 0
0
B A
B
B A
B
B A
B
B
B
dt = IJ (_) - 2 = 4
B A
d 2 =5
dt = It(_) - 1 = 5
d=max(4, 5) =5
d2 =2
d = max(5, 2) = 5
B A
0
FIGURE 7.3 Example of string matching with the Boyer-Moore algorithm
The good-suffix table is filled as follows:
k
pattern
dz
1
2
3
4
5
BAOBA~
BAOBAB
BAOBAB
BAOBAB
BAOBAB
2
5
5
5
5
The actual search for this pattern in the text given in Figure 7.3 proceeds as
follows. After the last Bof the pattern fails to match its counterpart Kin the text, the
algorithm retrieves t 1(K) = 6 from the bad-symbol table and shifts the pattern by
dt = max(lt (K) - 0, 1) = 6 positions to the right. The new try successfully matches
two pairs of characters. After the failure of the third comparison on the space
character in the text, the algorithm retrieves t1(_) = 6 from the bad-symbol table
and d2 = 5 £rom the good-suffix table to shift the pattern bymax(d1,d2 ) = max(62, 5) = 5. Note that on this iteration it is the good-suffix rule that leads to a farther
shift of the pattern.
The next try successfully matches just one pair of B's. After the failure of the
next comparison on the space character in the text, the algorithm retrieves It(_)
= 6 from the bad-symbol table and d 2 = 2 from the good-suffix table to shift the
pattern by max(dt,dz) = max(6- 1, 2) = 5. Note that on this iteration it is the badsymbol rule that leads to a farther shift of the pattern. The next try finds a matching
substring in the text after successfully matching all six characters of the pattern
with their counterparts in the text.
Ill
When searching for the first occurrence of the pattern, the worst-case efficiency of the Boyer-Moore algorithm is known to be linear. Though this algorithm
s
~·. ·.~ ·.!·
"
•·. .•·•
! !
.
_:,j
264
Space and Time Tradeoffs
runs very fast, especially on large alphabets (relative to the length of the pattern),
many people prefer its simplified versions, such as Horspool's algorithm, when
dealing with natural language-like strings.
-----Exercises 7 . 2 - - - - - - - - - - - - - - - 1. Apply Horspool's algorithm to search for the pattern BAOBAB in the text
BESS_KNEW_ABOUT_BAOBABS
2. Consider the problem of searching for genes in DNA sequences using Horspool's algorithm. A DNA sequence is represented by a text on the alphabet
{A, C, G, T), and the gene or gene segment is the pattern.
a. Construct the shift table for the following gene segment of your chromosome 10:
TCCTATICTI
b. Apply Horspool's algorithm to locate the above pattern in the following
DNA sequence:
TIATAGATCTCGTATICTITIATAGATCTCCTATICTI
3. How many character comparisons will be made by Horspool's algorithm in
searching for each of the following patterns in the binary text of 1000 zeros?
a. 00001
b. 10000
c. 01010
4. For searching in a text of length n for a pattern of length m (n 2: m) with
Horspool's algorithm, give an example of
b. best-case input.
a. worst-case input.
5. Is it possible for Horspool's algorithm to make more character comparisons
than the brute-force algorithm would make in searching for the same pattern
in the same text?
6. If Horspool's algorithm discovers a matching substring, how large a shift
should it make to search for a uext possible match?
7. How many character comparisons will the Boyer-Moore algorithm make in
searching for each of the following patterns in the binary text of 1000 zeros?
a. 00001
b. 10000
c. 01010
8. a. Would the Boyer-Moore algorithm work correctly with just the bad-symbol
table to guide pattern shifts?
b. Would the Boyer-Moore algorithm work correctly with just the good-suffix
table to guide pattern shifts?
7.3
Hashing
265
9. a. If the last characters of a pattern and its counterpart in the text do match,
does Horspool"s algorithm have to check other characters right to left, or
can it check them left to right too?
b. Answer the same question for the Boyer-Moore algorithm.
10. Implement Horspool"s algorithm, the Boyer-Moore algorithm, and the bruteforce algorithm of Section 3.2 in the language of your choice and run an
experiment to compare their efficiencies for matching
a. random binary patterns in random binary texts.
b. random natural language patterns in naturallauguage texts.
7.3 Hashing
In this section, we consider a very efficient way to implement dictionaries. Recall
that a dictionary is an abstract data type, namely, a set with the operations of
searching, insertion, and deletion defined on its elements. The elements of this set
can be of an arbitrary nature: numbers, characters of some alphabet, character
strings, and so on. In practice, the most important case is that of records (student
records in a school, citizen records in a governmental office, book records in a
library).
Typically, records comprise several fields, each responsible for keeping a
particular type of information about an entity the record represents. For example,
a student record may contain fields for the student's ID, name, date of birth, sex,
home address, major, and so on. Among record fields there is usually at least one
called a key that is used for identifying entities represented by the records (e.g., the
student's ID). In the following discussion, we assume that we have to implement
a dictionary of n records with keys K1, K 2 , •.. , Kn.
Hashing is based on the idea of distributing keys among a one-dimensional
array H[O .. m- 1] called a hash table. The distribution is done by computing,
for each of the keys, the value of some predefined function h called the hash
function. This function assigns an integer between 0 and m - 1, called the hash
address, to a key. For example, if keys are nonnegative integers, a hash function
can be of the form h(K) = K mod m (obviously, the remainder of division by m
is always between 0 and m - 1). If keys are letters of some alphabet, we can first
assign a letter its position in the alphabet (denoted here ord(K)) and then apply
the same kind of a function used for integers. Finally, if K is a character string
coc1 ... c,_lo we can use, as a very unsophisticated option,
ord(c;)) mod m;
a better option is to compute h(K) as follows: 2
CL:;:t
h <-- 0; fori <-- 0 to s- 1 doh <-- (h
2.
* C + ord(c;)) mod m,
'This can be obtained by treating ord(ci) as digits of a number in the C-hased system, computing its
decimal value by Horner's rule, and finding the remainder of the number after dividing it by m.
Ll
266
Space and Time Tradeoffs
FIGURE 7.4 Collision of two keys in hashing: h(K;) = h(K;)
where Cis a constant larger than every ord(c;).
In general, a hash function needs to satisfy two somewhat conflicting requirements:
1!!1
1!!1
A hash function needs t<Mii'"trib\]teJ.~y~among the cells of the hash table
as evenly as possible. (Because of this requirement, the value of m is usually
chosen to be prime. This requirement also makes it desirable, for most applications, to have a hash function dependent on all bits of a key, not just some
of them.)
A hash function has to be.eas-y•to compute.
Obviously, if we choose a hash table's size m to be smaller than the number of
keys n, we will get collisions-a phenomenon of two (or more) keys being hashed
into the same cell of the hash table (Figure 7.4). But collisions should be expected
even if m is considerably larger than n (see Problem 5). In fact, in the worst case,
all the keys could be hashed to the same cell of the hash table. Fortunately, with an
appropriately chosen size of the hash table and a good hash function, this situation
will happen very rarely. Still, every hashing scheme must have a collision resolution
mechanism. This mechanism is different in the two principal versions of hashing:
open hashing (also called separate chaining) and closed hashing (also called open
addressing).
Open Hashing (Separate Chaining)
In open hashing, keys are stored in linked lists attached to cells of a hash table.
Each list contains all the keys hashed to its cell. Consider, as an example, the
following list of words:
A,FOOL,AND,HIS,MONEY,ARE,SOON,PARTED.
As a hash function, we will use the simple function for strings mentioned above,
that is, we will add the positions of a word's letters in the alphabet and compute
the sum's remainder after division by 13.
We start with the empty table. The first key is the word A; its hash value is
h(A) = 1 mod 13 = 1. The second key-the word FOOL-isinstalled in the ninth cell
(since (6 + 15 + 15 + 12) mod 13 = 9), and so on. The final result of this process is
given in Figure 7.5; note a collision of the keys ARE and SOON (because h(ARE) =
(1 + 18 + 5) mod 13 = 11 and h(SOON) = (19 + 15 + 15 + 14) mod 13 = 11).
7.3
Hashing
9
10
11
12
FOOL
HIS
ARE
PARTED
267
keys
hash addresses
0
2
A
3
4
5
6
7
AND
MONEY
8
J.
SOON
FIGURE 7.5 Example of a hash table construction with separate chaining
How do we search in a dictionary implemented as such a table of linked lists?
We do this by simply applying to a search key the same procedure that was used
for creating the table. To illustrate, if we want to search for the key KID in the hash
table of Figure 7.5, we first compute the value of the same hash function for the
key: h(KID) = 11. Since the list attached to cellll is not empty, its linked list may
contain the search key. But because of possible collisions, we cannot tell whether
this is the case until we traverse this linked list. After comparing the string KID first
with the string ARE and then with the string SOON, we end up with an unsuccessful
search.
In general, the efficiency of searching depends on the lengths of the liuked
lists, which, in turn, depend on the dictionary and table sizes, as well as the quality
of the hash function. If the hash function distributes n keys among m cells of the
hash table about evenly, each list will be about n/m keys long. The ratio a= n/m,
called the load factor of the hash table, plays a crucial role in the efficiency of
hashing. In particular, the average number of pointers (chain links) inspected in
successful searches, S, and unsuccessful searches, U, turn out to be
S"'1+
"z and U=a,
(7,4)
respectively (under the standard assumptions of searching for a randomly selected
element and a hash function distributing keys uniformly among the table's cells).
These results are quite natural. Indeed, they are almost identical to searching
sequentially in a linked list; what we have gained by hashing is a reduction in
average list size by a factor of m, the size of the hash table.
Normally, we want the load factor to be not far from 1. Having it too small
would imply a lot of empty lists and hence inefficient use of space; having it too
large would mean longer linked lists and hence longer search times. But if the load
factor is around 1, we have an amazingly efficient scheme that makes it possible to
search for a given key for, on average, the price of one or two comparisons! True,
in addition to comparisons, we need to spend time on computing the value of the
~·n·I,~. ~
f
·
..·.••.......... •.•
..·.:
..
268
Space and Time Tradeoffs
hash function for a search key, but it is a constant-time operation, independent
from n and m. Note that we are getting this remarkable efficiency not only as a
result of the method's ingenuity but also at the expense of extra space.
The two other dictionary operations-insertion and deletion-are almost
identical to searching. Insertions are normally done at the end of a list (but see
Problem 6 for a possible modification of this rule). Deletion is performed by
searching for a key to be deleted and then removing it from its list. Hence, the
efficiency of these operations is identical to that of searching, and they are all
8(1) in the average case if the number of keys n is about equal to the hash table's
size m.
Closed Hashing (Open Addressing)
In closed hashing, all keys are stored in the hash table itself without the use
of linked lists. (Of course, this implies that the table size m must be at least as
large as the number of keys n.) Different strategies can be employed for collision
resolution. The simplest one-called linear probing-checks the cell following the
one where the collision occurs. If that cell is empty, the new key is installed there; if
the next cell is already occupied, the availability of that cell's immediate successor
is checked, and so on. Note that if the end of the hash table is reached, the search
is wrapped to the beginning of the table; that is, it is treated as a circular array.
This method is illustrated in Figure 7.6 by applying it to the same list of words we
used above to illustrate separate chaining (we also use the same hash function).
To search for a given key K, we start by computing h(K) where his the hash
function used in the table's construction. If the cell h(K) is empty, the search is
unsuccessful. If the cell is not empty, we must compare K with the cell's occupant:
keys
hash addresses
2
0
PARTED
A
A
A
A
A
A
A
A
3
4
5
6
AND
AND
AND
AND
AND
AND
7
MONEY
MONEY
MONEY
MONEY
8
9
10
11
12
FOOL
FOOL
FOOL
FOOL
FOOL
FOOL
FOOL
HIS
HIS
HIS
HIS
HIS
ARE
ARE
ARE
SOON
SOON
FIGURE 7.6 Example of a hash table construction with linear probing
,z···
7.3
Hashing
269
if they are equal, we have found a matching key; if they are not, we compare
K with a key in the next cell and continue in this manner until we encounter
either a matching key (a successful search) or an empty cell (unsuccessful search).
For example, if we search for the word LIT in the table of Figure 7.6, we will
get h (LIT) = (12 + 9 + 20) mod 13 = 2 and, since cell 2 is empty, we can stop
immediately. However, if we search for KID with h(KID) = (11 + 9 + 4) mod 13 =
11, we will have to compare KID with ARE, SOON, PARTED, and A before we can
declare the search unsuccessful.
While the search and insertion operations are straightforward for this version
of hashing, deletion is not. For example, if we simply delete the key ARE from the
last state of the hash table in Figure 7.6, we will be unable to find the key SOON
afterward. Indeed, after computing h(SOON) = 11, the algorithm would find this
location empty and report the unsuccessful search result. A simple solution is to
use "lazy deletion," that is, to mark previously occupied locations by a special
symbol to distinguish them from locations that have not been occupied.
The mathematical analysis of linear probing is a much more difficult problem
than that of separate chaining 3 The simplified versions of these results state that
the average number of times the algorithm must access a hash table with the load
factor a in successful and unsuccessful searches is, respectively,
1
1
]
S"'-(1+--) and U"'-(1+
2
1- a
2
]
(1 - a )2
)
(7.5)
(and the accuracy of these approximations increases with larger sizes of the hash
table). These numbers are surprisingly small even for densely populated tables,
i.e., for large percentage values of a:
1(1 + _1_)
1-a
2
50%
1.5
75%
2.5
90%
5.5
2.5
8.5
50.5
Still, as the hash table gets closer to being full, the performance of linear probing deteriorates because of a phenomenon called clustering. A cluster in linear
probing is a sequence of contiguously occupied cells (with a possible wrapping).
For example, the final state of the hash table in Figure 7.6 has two clusters. Clusters
are bad news in hashing because they make the dictionary operations less efficient. Also note that as clusters become larger, the probability that a new element
3.
This problem was solved in 1962 by a young graduate student in mathematics named Donald E. Knuth.
Knuth went on to become one of the most important computer scientists of our time. His multivolume
treatise The Art of Computer Programming [Knul, Knuli, Knuiii] remains the most comprehensive
and influential book on algorithmics ever published.
270
Space and 1ime Tradeoffs
is attached to a cluster increases; in addition, large clusters increase the probability that two clusters will coalesce after a new key's insertion, causing even more
clustering.
Several other collision resolution strategies have been suggested to alleviate
this problem. One of the most important is double hashing. Under this scheme, we
use another hash function, s(K), to determine a fixed increment for the probing
sequence to be used after a collision at location l = h(K):
(I+ s(K)) mod m, (l + 2s(K)) mod m, ... .
(7.6)
For every location in the table to be probed with sequence (7.6), the increment
s(K) and the table size m must be relatively prime, i.e., their only common divisor
must be l. (This condition is satisfied automatically if m itself is prime.) Some
functions recommended in the literature are s(k) = m- 2- k mod (m- 2) and
s(k) = 8- (k mod 8) for small tables and s(k) = k mod 97 + 1 for larger ones (see
[Sed88], [Sed98]). Mathematical analysis of double hashing has proved to be
quite difficult. Some partial results and considerable practical experience with the
method suggest that with good hashing functions-both primary and secondarydouble hashing is superior to linear probing. But its performance also deteriorates
when the table gets close to being full. A natural solution in such a situation is
rehashing: the current table is scanned, and all its keys are relocated into a larger
table.
Since its discovery in the 1950s by IBM researchers, hashing has found many
important applications. In particular, it has become a standard technique for storing a symbol table-a table of a computer program's symbols generated during
compilation. With some modifications, it has also proved to be useful for storing
very large dictionaries on disks; this variation of hashing is called extendible hashing. Since disk access is expensive compared with probes performed in the main
memory, it is preferable to make many more probes than disk accesses. Accordingly, a location computed by a hash function in extendible hashing indicates a
disk address of a bucket that can hold up to b keys. When a key's bucket is identified, all its keys are read into main memory and then searched for the key in
question. In the next section, we discuss B-trees, a principal alternative for storing
large dictionaries.
------Exercises
1. For the input 30, 20, 56, 75, 31, 19 and hash function h(K) = K mod 11
a. construct the open hash table.
b. find the largest number of key comparisons in a successful search in this
table.
c. find the average number of key comparisons in a successful search in this
table.
2. For the input 30, 20, 56, 75, 31, 19 and hash function h(K) = K mod 11
7.3
Hashing
271
a. construct the closed hash table.
b. find the largest number of key comparisons in a successful search in this
table.
c. find the average number of key comparisons in a successful search in this
table.
(-·--~,
i 3.; Why is it not a good idea for a hash function to depend on just one letter (say,
I t the first one) of a natural language word?
l
i
i 4.\ Find the probability of all n keys being hashed to the same cell of a hash table
L ..J of size m if the hash function distributes keys evenly among all the cells of the
table.
5. Birthday paradox The birthday paradox asks how many people should be
in a room so that the chances are better than even that two of them will have
the same birthday (month and day). Find the quite unexpected answer to this
problem. What implication for hashing does this result have?
6. Answer the following questions for the separate-chaining version of hashing.
a. Where would you insert keys if you knew that all the keys in the dictionary
are distinct? Which dictionary operations, if any, would benefit from this
modification?
b. We could keep keys of the same linked list sorted. Which of the dictionary operations would benefit from this modification? How could we take
advantage of this if all the keys stored in the entire table need to be sorted?
(7;! Explain how hashing can be applied to check whether all elements of a list
·-...•;·are distinct. What is the time efficiency of this application?
8. Fill in the following table with the average-case efficiency classes for the five
implementations of the ADT dictionary:
unordered
array
ordered
array
binary
separate
search tree
chaining
linear
probing
search
insertion
deletion
9. We have discussed hashing in the context of techniques based on space-time
tradeoffs. But it also takes advantage of another general strategy. Which one?
10. Write a computer program that uses hashing for the following problem. Given
a natural language text, generate a list of distinct words with the number of
occurrences of each word in the text. Insert appropriate counters in the program to compare the empirical efficiency of hashing with the corresponding
theoretical results.
272
Space and Time Tradeoffs
7.4 B-Trees
The idea of using extra space to facilitate faster access to a given data set is particularly important if the data set in question contains a very large number of records
that need to be stored on a disk. A principal device in organizing such data sets
is an index, which provides some information about the location of records with
indicated key values. For data sets of structured records (as opposed to "unstructured" data such as text, images, sound, and video), the most important index
organization is the B-tree, introduced by R. Bayer and E. McGreight [Bay72]. It
extends the idea of the 2-3 tree (see Section 6.3) by permitting more than a single
key in the same node of a search tree.
In the B-tree version we consider here, all data records (or record keys)
are stored at the leaves, in increasing order of the keys. The parental nodes are
used for indexing. Specifically, each parental node contains n - 1 ordered keys
K1 < ... < K,_ 1 assumed, for the sake of simplicity, to be distinct. The keys are
interposed with n pointers to the node's children so that all the keys in subtree
To are smaller than K l> all the keys in subtree T1 are greater than or equal to
K1 and smaller than K 2 with K 1 being equal to the smallest key in Th and so on
through the last subtree T,,_ 1 whose keys are greater than or equal than K,_ 1
with K,_ 1 being equal to the smallest key in T,,_ 1 (Figure 7.7). The node depicted
in Figure 7.7 is called the n-node. Thus, all the nodes in a classic binary search
tree are 2-nodes; a 2-3 tree introduced in Section 6.3 comprises 2-nodes and 3nodes.
In addition, a B-tree of order m 2: 2 must satisfy the following structural
properties.
FIGURE 7. 7 Parental node of a B-tree
7.4
B-Trees
273
FIGURE 7.8 Example of a B-tree of order 4
"
Ill
Iii
The root is either a leaf or has between 2 and rn children.
Each node, except for the root and the leaves, has between frn /21 and rn
children (and hence between fm/21 -1 and rn -1 keys).
The tree is (perfectly) balanced; i.e., all its leaves are at the same level.
An example of a B-tree of order 4 is given in Figure 7.8.
Searching in a B-tree is very similar to searching in the binary search tree,
and even more so in the 2-3 tree. Starting at the root, we follow a chain of pointers
to the leaf that may contain the search key. Then we search for the search key
among the keys of that leaf. Note that since keys are stored in sorted order, at
both parental nodes and leaves, we can use binary search if the number of keys at
a node is large enough to make it worthwhile.
It is not the number of key comparisons, however, that we should be concerned about in a typical application of this data structure. When used for storing
a large data file on a disk, the nodes of a B-tree normally correspond to the disk
pages. Since the time needed to access a disk page is typically several orders of
magnitude larger than the time needed to compare keys in the fast computer memory, it is the number of disk accesses that becomes the principal indicator of the
efficiency of this and similar data structures.
How many nodes of a B-tree do we need to access during a search for a record
with a given key value? This number is obviously equal to the height of the tree
plus one. To estimate the height, let us find the smallest number of keys a B-tree
of order rn and positive height h can have. The root of the tree will contain at least
one key. Levell will have at least two nodes with at least fm /21 - 1 keys in each
of them, for the total minimum number of keys 2(frn/21 - 1). Level2 will have at
least 2 rm /21 nodes (the children of the nodes on level 1) with at least rm /21 - 1
in each of them, for the total minimum number of keys 2fm/2l(fm/21- 1). In
general, the nodes of level i, 1 ""'i ""'h- 1, will contain at least 2frn/21'-· 1(fm/21 1) keys. Finally, level h, the leaf level, will have at least 2fm/21"- 1 nodes with at
~.·
·r
!
•.'•.·••·••.••••••
••.· • ·..1•
••• ·•·•
•
f"'
274
Space and lime Tradeoffs
!
least one key in each. Thus, for any B-tree of order m with n nodes and height
h > 0, we have the following inequality:
h~l
n 2:1 +
L 2rm/2l'- Um/2l -1) +2rm/2lh-l.
1
i=l
After a series of standard simplifications (Problem 2), this inequality reduces to
n 2: 4rm/2lh-l -1,
which, in turn, yields the following upper bound on the height h of the B-tree of
order m with n nodes:
n+1
(7.7)
h ::0 [logfm/21 -4-J+ 1.
Inequality (7.7) immediately implies that searching in a B-tree is a O(log n)
operation. But it is important to ascertain here not just the efficiency class but
the actual number of disk accesses implied by this formula. The following table
contains the values of the right-hand side estimates for a file of 100 million records
and a few typical values of the tree's order m:
orderm
h's upper bound
50
100
250
6
5
4
Keep in mind that the table's entries are upper estimates for the number of disk
accesses. In actual applications, this number rarely exceeds 3, with the B-tree's
root and sometimes first-level nodes stored in the fast memory to minimize the
number of disk accesses.
The operations of insertion and deletion are less straightforward than searching, but both can also be done in 0 (log n) time. Here we outline an insertion
algorithm only; a deletion algorithm can be found in the references (e.g., [Aho83],
[CorOl]).
The most straightforward algorithm for inserting a new record into a B-tree
is quite similar to the algorithm for insertion into a 2-3 tree outlined in Section
6.3. First, we apply the search procedure to the new record's key K to find the
appropriate leaf for the new record. If there is room for the record in that leaf,
we place it there (in an appropriate position so that the keys remain sorted) and
we are done. If there is no room for the record, the leaf is split in half by sending
the second half of the records to a new node. After that, the smallest key K' in
the new node and the pointer to it have to be inserted into the old leaf's parent
(inlmediately after the key and pointer to the old leaf). This recursive procedure
may percolate up to the tree's root. If the root is already full too, a new root is
created with the two halves of the old root's keys split between two children of
7.4
B-Trees
275
FIGURE 7.9 B-tree obtained after inserting 65 into the B-tree in Figure 7.8
the new root. As an example, Figure 7.9 shows the result of inserting 65 into the Btrce in Figure 7.8 under the restriction that the tree's leaves cannot contain more
than three items.
You should be aware that there are other algoritluns for implementing insertions into a B-tree. For example, we can avoid the possibility of recursive node
splits if we split full nodes as we are searching for an appropriate leaf for the new
record. Another possibility is to avoid some node splits by moving a key to the
node's sibling. For example, inserting 65 into the B-tree in Figure 7.8 can be done
by moving 60, the smallest key of the full leaf, to its sibling with keys 51 and 55, and
replacing the key value of their parent by 65, the new smallest value in the second
child. This modification tends to save some space at the expense of a slightly more
complicated algorithm.
A B-tree does not have to be always associated with the indexing of a large
file, and it can be considered as one of several varieties of the search tree. As
with other types of search trees-such as binary search trees, AVL trees, and 2-3
trees-a B-tree can be constructed by successive insertions of data records into the
initially empty tree. (The empty tree is considered to be a B-tree, too.) When all
keys reside in the leaves and the upper levels are organized as a B-tree comprising
an index, the entire structure is usually called, in fact, an+ -tree.
-----Exercises 7 . 4 - - - - - - - - - - - - - - - 1. Give examples of using an index in real-life applications that do not involve
computers.
2. a. Prove the equality
h-1
1 + I:2rm/21;-tcrm/21 -1) + 2rm/21h-l = 4rm/21h-l_l,
i=l
276
Space and Time Tradeoffs
which was used in the derivation of upper bound (7.7) for the height of a
B-tree.
b. Complete the derivation of inequality (7.7).
J<i
nuuthe minimum order of the B-tree that guarantees that the number of disk
accesses in searching in a file of 100 million records does not exceed 3. Assume
that the root's page is stored in main memory.
4. Draw the B-tree obtained after inserting 30 and then 31 in the B-tree in
Figure 7.8. Assume that a leaf cannot contain more than three items.
5. Outline an algorithm for finding the largest key in a B-tree.
6. a. A top-down 2-3-4 tree is a B-tree of order 4 with the following modification of the insert operation. Whenever a search for a leaf for a new key
encounters a full node (i.e., a node with three keys), the node is split into
two nodes by sending its middle key to the node's parent (or, if the full
node happens to be the root, the new root for the middle key is created).
Construct a top-down 2-3-4 tree by inserting tbe following list of keys in
the initially empty tree:
10,6, 15,31,20,27,50,44, 18.
b. What is the principal advantage of this insertion procedure compared with
the one used for 2-3 trees in Section 6.3? What is its disadvantage?
7. a. Write a program implementing a key insertion algorithm in a B-tree.
b. Write a program for visualization of a key insertion algorithm in a B-tree.
SUMMARY
'"
Space and time tradeoffs in algorithm design are a well-known issue for
both theoreticians and practitioners of computing. As an algorithm design
technique, trading space for time is much more prevalent than trading time
for space.
11
Input enhancement is one of the two principal varieties of trading space for
time in algorithm design. Its idea is to preprocess the problem's input, in whole
or in part, and store the additional information obtained in order to accelerate
solving the problem afterward. Sorting by distribution counting and several
important algorithms for string matching are examples of algorithms based
on this technique.
01
Distribution counting is a special method for sorting lists of elements from a
small set of possible values.
Summary
277
"
Horspool:v algorithm for string matching can be considered a simplified
version of the Boyer-Moore algorithm. Both algorithms are based on the ideas
of inpnt enhancement and right-to-left comparisons of a pattern's characters.
Both algorithms use the same bad-symbol shift table; the Boyer-Moore also
uses a second table, called the good-suffix shift table.
ill
?restructuring-the second type of technique that exploits space-for-time
tradeoffs-uses extra space to facilitate a faster and/or more flexible access
to the data. Hashing and B-trees are important examples of prestructuring.
"
Hashing is a very efficient approach to implementing dictionaries. It is based
on the idea of mapping keys into a one-dimensional table. The size limitations
of such a table make it necessary to employ a collision resolution mechanism.
The two principal varieties of hashing are open hashing or separate chaining
(with keys stored in linked lists outside of the hash table) and closed hashing
or open addressing (with keys stored inside the table). Both enable searching,
insertion, and deletion in 8(1) time, on average.
"'
Tiw B-tree is a balanced search tree that generalizes the idea of the 2-3 tree
by allowing multiple keys at the same node. Its principal application is for
keeping index-like information about data stored on a disk. By choosing
the order of the tree appropriately, one can implement the operations of
searching, insertion, and deletion with just a few disk accesses even for
extremely large files.
r
!
I!
-..,
!
Dynamic Programming
An idea, like a ghost ... must be spoken to a little before it will explain itself.
-Charles Dickens 11812-1870)
D
ynamic programming is an algorithm design technique with a rather interesting history. It was invented by a prominent U.S. mathematician, Richard
Bellman, in the 1950s as a general method for optimizing multistage decision processes. Thus, the word "programming" in the name of this technique stands for
"planning" and does not refer to computer programming. After proving its worth
as an important tool of applied mathematics, dynamic programming has eventually come to be considered, at least in computer science circles, as a general
algorithm design technique that does not have to be limited to special types of
optimization problems. It is from this point of view that we will consider this technique here.
Dynamic programming is a technique for solving problems with overlapping
subproblems. Typically, these subproblems arise from a recurrence relating a
solution to a given problem with solutions to its smaller subproblems of the same
type. Rather than solving overlapping subproblems again and again, dynamic
programming suggests solving each of the smaller subproblems only once and
recording the results in a table from which we can then obtain a solution to the
original problem.
This technique can be illustrated by revisiting the Fibonacci numbers discussed in Section 2.5. (If you have not read that section, you will be able to follow
the discussion anyway. But it is a beautiful topic, so if you feel a temptation to read
it, do succumb to it.) The Fibonacci numbers are the elements of the sequence
0, 1, 1, 2, 3, 5, 8, 13, 21, 34, ... ,
which can be defined by the simple recurrence
F(n) = F(n- 1) + F(n- 2)
for n )> 2
(8.1)
and two initial conditions
F(O) = 0, F(1) = 1.
(8.2)
279
I
280
Dynamic Programming
If we try to use recurrence (8.1) directly to compute the nth Fibonacci number
F (n), we would have to recompute the same values of this function many times
(see Figure 2.6 for a specific example). Note that the problem of computing F(n)
is expressed in terms of its smaller and overlapping subproblems of computing
F(n- 1) and F(n- 2). So we can simply fill elements of a one-dimensional array
with then+ 1 consecutive values of F(n) by starting, in view of initial conditions
(8.2), with 0 and 1 and using equation (8.1) as the rule for producing all the other
elements. Obviously, the last element of this array will contain F (n ). A single-loop
pseudocode for this very simple algorithm can be found in Section 2.5.
Note that we can, in fact, avoid using an extra array to accomplish this task
by recording the values of just the last two elements of the Fibonacci sequence
(see Problem 6 in Exercises 2.5). This phenomenon is not unusual, and we shall
encounter it in a few more examples in this chapter. Thus, although a straightforward application of dynamic programming can be interpreted as a special variety
of space-for-time tradeoff, a dynamic programming algorithm can sometimes be
refined to avoid using extra space.
Certain algorithms compute the nth Fibonacci number without computing
all the preceding elements of this sequence (see Section 2.5). It is typical of an
algorithm based on the classic bottom-up dynamic programming approach, however, to solve all smaller subproblems of a given problem. One variation of the
dynamic programming approach seeks to avoid solving unnecessary subproblems.
This technique, illustrated in Section 8.4, exploits so-called memory functions and
can be considered a top-down variation of dynamic programming.
Whether we use the classical bottom-up version of dynamic programming or
its top-down variation, the crucial step in designing such an algorithm remains the
same: namely, deriving a recurrence relating a solution to the problem's instance to
solutions of its smaller (and overlapping) subinstances. The immediate availability
of equation (8.1) for computing the nth Fibonacci number is one of the few
exceptions to this rule.
In the sections and exercises of this chapter are a few standard examples of
dynamic programming algorithms. (Some of them, in fact, were invented before or
independent of the discovery of dynamic programming and only later came to be
viewed as examples of this technique's applications.) Numerous other applications
range from the optimal way of breaking text into lines (e.g., [BaaOO]) to an
optimal triangulation of a polygon (e.g., [Ski98]) to a variety of applications to
sophisticated engineering problems (e.g., [Bel62], [Ber01]).
8.1
Computing a Binomial Coefficient
Computing a binomial coefficient is a standard example of applying dynamic
programming to a nonoptimization problem. You may recall from your studies
of elementary combinatorics that the binomial coefficient, denoted C(n, k) or
G), is the number of combinations (subsets) of k elements from an n-element
i
8.1
Computing a Binomial Coefficient
281
set (0 :<= k :<= n). The name "binomial coefficients" comes from the participation of
these numbers in the binomial formula:
(a+ b)"= C(n, O)a" +. · · + C(n, k)a"-kbk + · · · + C(n, n)b".
Of the numerous properties of binomial coefficients, we concentrate on two:
C(n, k) = C(n- 1, k- 1) + C(n- 1, k)
for n > k > 0
(8.3)
and
C(n, 0) = C(n, n) = 1.
(8.4)
The nature of recurrence (8.3), which expresses the problem of computing C(n, k) in terms of the smaller and overlapping problems of computing
C (n - 1, k - 1) and C (n - 1, k), lends itself to solving by the dynamic programming technique. To do this, we record the values of the binomial coefficients in
a table of n + 1 rows and k + 1 columns, numbered from 0 to n and from 0 to k,
respectively (Figure 8.1).
To compute C(n, k), we fill the table in Figure 8.1 row by row, starting with
row 0 and ending with row n. Each row i (0 :<= i :<= n) is filled left to right, starting
with 1 because C(n, 0) = 1. Rows 0 through k also end with 1 on the table's
main diagonal: C(i, i) = 1 for 0 :<= i :<= k. We compute the other entries by formula
(8.3), adding the contents of the cells in the preceding row and the previous
column and in the preceding row and the same column. (If you recognize Pascal's
triangle-a fascinating mathematical structure usually studied in conjunction with
the notion of a combination-you are right: this is exactly what it is.) The following
pseudocode implements this algorithm.
0
2
...
k-1
k
0
2
2
k
n-1
n
C(n-1, k-1)
C(n-1, k)
C(n, k)
FIGURE 8.1 Table for computing the binomial coefficient C(n, k) by the dynamic
programming algorithm
:;:::
282
Dynamic Programming
ALGORITHM Binomial(n, k)
//Computes C(n, k) by the dynamic programming algorithm
//Input: A pair of nonnegative integers n ": k ": 0
//Output: The value of C(n, k)
for i <--- 0 to n do
for j <--- 0 to min(i, k) do
if j = 0 or j = i
C[i, j] <---1
else C[i, J] <--- C[i -1, j -1] + C[i -1, j]
return C[n, k]
What is the time efficiency of this algorithm? Clearly, the algorithm's basic
operation is addition, so let A (n, k) be the total number of additions made by this
algorithm in computing C(n, k). Note that computing each entry by formula (8.3)
requires just one addition. Also note that because the first k + 1 rows of the table
form a triangle while the remaining n - k rows form a rectangle, we have to split
the sum expressing A (n, k) into two parts:
k
A(n, k)
i-1
k
k
i=k+l j=l
i=1
11
II
=I: 2:) + I: I: 1 =I;u- 1) + I: k
i=l j=l
=
(k- 1)k
2
+ k(n- k) E EJ(nk).
You are asked to ascertain whether this is an efficient algorithm by comparing it
with the running times of a few other algorithms for this problem in the exercises.
Another problem in the exercises is to analyze whether or not the extra space used
by the dynamic programming algorithm is actually necessary.
-----Exercises 8.1 - - - - - - - - - - - - - - 1. a. What does dynamic programming have in common with divide-andconquer?
b. What is a principal difference between the two techniques?
2. a. Compute C(6, 3) by applying the dynamic programming algorithm.
b. Is it also possible to compute C(n, k) by filling the algorithm's dynamic
programming table column by column rather than row by row?
3. Prove the following assertion made in the text while investigating the time
efficiency of the dynamic programming algorithm for computing C(n, k):
(k -1)k
2
+ k(n- k) E EJ(nk).
8.1
Computing a Binomial Coefficient
283
4. a. What is the space efficiency of Binomial, the dynamic programming algorithm for computing C(n, k)?
b. Explain how the space efficiency of this algorithm can be improved. (Try
to make as much of an improvement as you can.)
5. a. Find the order of growth of the following functions.
i. C(n, 1)
ii. C(n, 2)
iii. C(n, n/2) for even n's
b. What major implication for computing C(n, k) do the answers to the questions in part (a) have?
6. Find the exact number of additions made by the following recursive algorithm
based directly on formulas (8.3) and (8.4).
ALGORITHM BinomCoeff(n, k)
if k = 0 or k = n return 1
else return BinomCoeff(n - 1, k - 1)+BinomCoeff(n - 1, k)
7. Which of the following algorithms for computing a binomial coefficient is most
efficient?
a. Use the formula
C n k =
n!
(,)
k!(n-k)!
b. Use the formula
C(
n,
k)
=
n(n - 1) ... (n - k + 1)
k!
.
c. Apply recursively the formula
C(n, k) = C(n- 1, k- 1) + C(n- 1, k)
C(n, 0) = C(n, n) = 1.
for n > k > 0,
d. Apply the dynamic programming algorithm.
8, Prove that
C(n, k) = C(n, n - k)
for n 0: k 0: 0
and explain how this formula can be utilized in computing C(n, k).
9, Shortest path counting A chess rook can move horizontally or vertically to
any square in the same row or in the same column of a chessboard. Find the
number of shortest paths by which a rook can move from one corner of a
chessboard to the diagonally opposite corner. [Gar78], p. 10
a. by a dynamic programming algorithm.
b. by using elementary combinatorics.
284
Dynamic Programming
10. World Series odds Consider two teams, A and B, playing a series of games
until one of the teams wins n games. Assume that the probability of A winning
a game is the same for each game and equal to p, and the probability of
A losing a game is q = 1- p. (Hence, there are no ties.) Let P(i, j) be the
probability of A winning the series if A needs i more games to win the series
and B needs j more games to win the series.
a. Set up a recurrence relation for P(i, j) that can be used by a dynamic
programming algorithm.
b. Find the probability of team A winning a seven-game series if the probability of the team winning a game is 0.4.
c. Write a pseudocode of the dynamic programming algorithm for solving
this problem and determine its time and space efficiencies.
--.,
tl
!
I!
li
8.2 Warshall's and Floyd's Algorithms
In this section, we look at two well-known algorithms: Warshall's algorithm for
computing the transitive closure of a directed graph and Floyd's algorithm for the
all-pairs shortest-paths problem. These algorithms are based on essentially the
same idea, which we can interpret as an application of the dynamic programming
technique.
Warshall's Algorithm
Recall that the adjacency matrix A= {au} of a directed graph is the boolean matrix
that has 1 in its ith row and jth column if and only if there is a directed edge from
the ith vertex to the jth vertex. We may also be interested in a matrix containing
the information about the existence of directed paths of arbitrary lengths between
vertices of a given graph.
DEFINITION The transitive closure of a directed graph with n vertices can be
defined as the n-by-n boolean matrix T = !tu}, in which the element in the ith row
(1 ::0 i ::0 n) and the jth column (1 ::0 j ::0 n) is 1 if there exists a nontrivial directed
path (i.e., a directed path of a positive length) from the ith vertex to the jth vertex;
otherwise, tij is 0.
An example of a digraph, its adjacency matrix, and its transitive closure is given
in Figure 8.2.
We can generate the transitive closure of a digraph with the help of depth-first
search or breadth-first search. Performing either traversal starting at the ith vertex
gives the information about the vertices reachable from the ith vertex and hence
the columns that contain ones in the ith row of the transitive closure. Thus, doing
such a traversal for every vertex as a starting point yields the transitive closure in
its entirety.
1:
8.2
285
Warshall's and Floyd's Algorithms
a
c
b
d
a
b
c
1
0
1
0
1
0
d
:sl T 1] T il
b
A= c
d
d
0
0
1
(a)
0
0
0
0
0
0
(b)
b
T= c
d
(c)
FIGURE 8.2 (a) Digraph. (b) Its adjacency matrix. (c) Its transitive closure.
Since this method traverses the same digraph several times, we should hope
that a better algorithm can be found. Indeed, such an algorithm exists. It is called
Warshall's algorithm after S. Warshall [War62]. Warshall's algorithm constructs
the transitive closure of a given digraph with n vertices through a series of n-by-n
boolean matrices:
(8.5)
Each of these matrices provides certain information about directed paths in the
digraph. Specifically, the element 1 in the ith row and jth column of matrix R(k)
(k = 0, 1, ... , n) is equal to 1 if and only if there exists a directed path (of a positive
length) from the ith vertex to the jth vertex with each intermediate vertex, if any,
numbered not higher thank. Thus, the series starts with Ri01, which does not allow
any intermediate vertices in its paths; hence, Ri01is nothing else but the adjacency
matrix of the digraph. (Recall that the adjacency matrix contains the information
about one-edge paths, i.e., paths with no intermediate vertices.) RID contains the
information about paths that can use the first vertex as intermediate; thus, with
more freedom, so to speak, it may contain more ones than Ri01. In general, each
subsequent matrix in series (8.5) has one more vertex to use as intermediate for
its paths than its predecessor and hence may, but does not have to, contain more
ones. The last matrix in the series, Ri"l, reflects paths that can use all n vertices of
the digraph as intermediate and hence is nothing else but the digraph's transitive
closure.
The central point of the algorithm is that we can compute all the elements of
each matrix R(k) from its immediate predecessor R(k-l) in series (8.5). Let r 1j 1,
the element in the ith row and jth column of matrix R(kJ, be equal to 1. This
means that there exists a path from the i th vertex v1 to the jth vertex vi with each
intermediate vertex numbered not higher thank:
rg
v1, a list of intermediate vertices each numbered not higher than k, vi.
(8.6)
Two situations regarding this path are possible. In the first, the list of its intermediate vertices does not contain the kth vertex. Then this path from v; to vi has
intermediate vertices numbered not higher thank - 1, and therefore r1j- 11 is equal
to 1 as well. The second possibility is that path (8.6) does contain the kth vertex vk
+
m.· .~!~.·
i·ll
..
286
Dynamic Programming
among the intermediate vertices. Without loss of generality, we may assume that
vk occurs only once in that list. (If it is not the case, we can create a new path from
vi to v j with this property by simply eliminating all the vertices between the first
and last occurrences of vk in it.) With this caveat, path (8.6) can be rewritten as
follows:
vi, vertices numbered.:::
k- 1, vk> vertices numbered::=:: k- 1, vj.
The first part of this representation means that there exists a path from v, to vk with
each intermediate vertex numbered not higher thank- 1 (hence r,~-l) = 1), and
the second part means that there exists a path from vk to v j with each intermediate
vertex numbered not higher thank- 1 (hence r~~-l) = 1).
What we have just proved is that if r,~) = 1, then either rg-l) = 1 or both
r,~-l) = 1 and r ~~-!) = 1. It is easy to see that the converse of this assertion is also
true. Thus, we have the following formula for generating the elements of matrix
R(k) from the elements of matrix Rlk-1):
r(k) = r(k-l) or (r<kk-l)and rk(k-l)) .
I}
I}
I
(8.7)
}
Formula (8.7) is at the heart of Warshall's algorithm. This formula implies
the following rule for generating elements of matrix R(k) from elements of matrix
Rlk-l), which is particularly convenient for applying Warshall's algorithm by hand:
,.
If an element rij is 1 in Rlk-l), it remains 1 in Rlk)
,.
If an element rij is 0 in R(k-l), it has to be changed to 1 in R(k) if and only if
the element in its row i and column k and the element in its column j and row
k are both 1's in R(k-l)_ (This rule is illustrated in Figure 8.3.)
As an example, the application ofWarshall's algorithm to the digraph in Figure 8.2
is shown in Figure 8.4.
j
k
~
Rlk-11 =
k
i
I
1
t
0
.
_., 1
'--
FIGURE 8.3 Rule tor changing zeros in Warshall's algorithm
j
k
8.2 Warshall's and Floyd's Algorithms
a
c
b
d
:sl l 1l
f' 1j
r
r i~
l iI
1
RIDI=
d
Rl11=
b
0
0
0
c
0
0
d
1
0
0
0
a
b
c
0
0
0
0
0
b
c
0
0
d
1 L...!_
a
b
c
Rl31=
b
d
Ones reflect the existence of paths
with intermediate vertices numbered
not higher than 2, i.e., a and b
(note two new paths);
boxed row and column are used for getting R!3l_
0
0
0
0
d
1
1L.!._
iJ
a
b
c
d
0
0
0
0
0
c
0
0
d
1
a
0
0
b
c
1
1
1
1
c
0
0
0
d
1
b
Rl41=
boxed row and column are used for getting R(1J.
Ones reflect the existence of paths
with intermediate vertices numbered
not higher than 1, i.e., just vertex a
(note a new path from dto b);
boxed row and column are used for getting R(2J.
c
b
Ones reflect the existence of paths
with no intermediate vertices
(R!DJ is just the adjacency matrix);
d
1'0""
Ri'l=
287
Ones reflect the existence of paths
with intermediate vertices numbered
not higher than 3, i.e., a, b, and c
(no new paths);
boxed row and column are used for getting R(4J.
d
Ones reflect the existence of paths
with intermediate vertices numbered
not higher than 4, i.e., a, b, c, and d
(note five new paths).
FIGURE 8.4 Application of Warshall's algorithm to the digraph shown. New ones are in
bold.
Here is a pseudocode of Warshall's algorithm.
ALGORITHM
Warshall(A[l..n, l..n])
//Implements Warshall's algorithm for computing the transitive closure
//Input: The adjacency matrix A of a digraph with n vertices
//Output: The transitive closure of the digraph
R(O) ~A
fork ~ 1 to n do
fori ~Hondo
for j ~ 1 to n do
R(kl[i, j] ~ R(k-l)[i, j] or (RCk-ll[i, k] and R(k-ll[k, j])
return R(n)
--,,,
I
288
Dynamic Programming
Several observations need to be made about Warshall's algorithm. First, it is
remarkably succinct, is it not? Still, its time efficiency is only in E>(n 3). In fact, for
sparse graphs represented by their adjacency lists, the traversal-based algorithm
mentioned at the beginning of this section has a better asymptotic efficiency
than Warshall's algorithm (why?). We can speed np the above implementation
of Warshall's algorithm for some inputs by restructuring its innermost loop (see
Problem 4 in the exercises). Another way to make the algorithm run faster is to
treat matrix rows as bit strings and employ the bitwise or operation available in
most modern computer languages.
As to the space efficiency of Warshall's algorithm, the situation is similar to
that of the two earlier examples in this chapter: computing a Fibonacci number and
computing a binomial coefficient. Although we used separate matrices for recording intermediate results of the algorithm, this is, in fact, unnecessary. (Problem 3 in
the exercises asks you to find a way of avoiding this wasteful use of the computer
memory.) Finally, we shall see below how the underlying idea of Warshall's algorithm can be applied to the more general problem of finding lengths of shortest
paths in weighted graphs.
Floyd's Algorithm for the All-Pairs Shortest-Paths Problem
Given a weighted connected graph (undirected or directed), the all-pairs shortestpaths problem asks to find the distances (the lengths of the shortest paths) from
each vertex to all other vertices. It is convenient to record the lengths of shortest
paths in an n-by-n matrix D called the distance matrix: the element diJ in the ith
row and the jth column of this matrix indicates the length of the shortest path
from the ith vertex to the jth vertex (1 :=:: i, j :=:: n ). For an example, see Figure 8.5.
We can generate the distance matrix with an algorithm that is very similar
to Warshall's algorithm. It is called Floyd's algorithm, after its inventor R. Floyd
[Flo62]. It is applicable to both undirected and directed weighted graphs provided
that they do not contain a cycle of a negative length. (Of course, in the case of a
directed graph, by a path or cycle we mean a directed path or a directed cycle.)
a b
c
d
a b
c
10
2 0
7 7
6 16
3
d
t8: T ~] T i]
3
7
1
(a)
w= b
2
c
=
d
6
0
7
0
D=
b
c
d
(b)
FIGURE 8.5 (a) Digraph. (b) Its weight matrix. (c) Its distance matrix.
5
0
9
(c)
8.2
Warshall's and Floyd's Algorithms
289
Floyd's algorithm computes the distance matrix of a weighted graph with n
vertices through a series of n-by-n matrices:
(8,8)
Each of these matrices contains the lengths of shortest paths with certain constraints on the paths considered for the matrix in question. Specifically, the element
d,~J in the ith row and the jth column of matrix n<k) (k = 0, 1, ... , n) is equal to
the length of the shortest path among all paths from the ith vertex to the jth vertex
with each intermediate vertex, if any, numbered not higher thank. In particular,
0
the series starts with n< l, which does not allow any intermediate vertices in its
paths; hence, n<OJ is nothing but the weight matrix of the graph. The last matrix
in the series, nrn), contains the lengths of the shortest paths among all paths that
can use all n vertices as intermediate and hence is nothing but the distance matrix
being sought.
As in Warshall's algorithm, we can compute all the elements of each matrix
n<'l from its immediate predecessor n<'- 1l in series (8.8). Let d,jl be the element
in the ith row and the jth column of matrix n<'l. This means that d,~J is equal to
the length of the shortest path among all paths from the ith vertex v; to the jth
vertex v1 with their intermediate vertices numbered not higher thank:
V;,
a list of intermediate vertices each numbered not higher than k, v .
1
(8.9)
We can partition all such paths into two disjoint subsets: those that do not use the
kth vertex as intermediate and those that do. Since the paths of the first subset
have their intermediate vertices numbered not higher thank- 1, the shortest of
1)
them is, by definition of our matrices, of length
What is the length of the shortest path in the second subset? If the graph does
not contain a cycle of a negative length, we can limit our attention only to the
paths in the second subset that use vertex as their intermediate vertex exactly
once (because visiting more than once can only increase the path's length). All
such paths have the following form:
v,
d,j-
v,
v,
vi, vertices numbered::; k -1, vk> vertices numbered::; k -1., v .
1
v, v,
In other words, each of the paths is made up of a path from
to
with each
intermediate vertex numbered not higher than k - 1 and a path from v, to v
with each intermediate vertex numbered not higher thank- 1. The situation is1
depicted symbolically in Figure 8.6.
Since the length of the shortest path from v; to v, among the paths that use
intermediate vertices numbered not higher than k - 1 is equal to d,~- 1 ) and the
length of the shortest path from v, to v1 among the paths that use intermediate
1
vertices numbered not higher thank - 1 is equal to
\ the length of the shortest
d;t
290
Dynamic Programming
d~.k-1)
'I
lk-1)
d kj
(k-1)
dlk
FIGURE 8.6 Underlying idea of Floyd's algorithm
path among the paths that use the kth vertex is equal to d1~- 1 ) + d~]- 1 ). Taking into
account the lengths of the shortest paths in both subsets leads to the following
recurrence:
d(k) = min{dlk- 1)
I)
I}
'
ik-t)
tk
1))
+ dlkk}
fork> 1
-
'
(8.10)
To put it another way, the element in the ith row and the jth column of the current
distance matrix Dlk- 1) is replaced by the sum of the elements in the same row i
and the kth column and in the same column j and the kth column if and only if
the latter sum is smaller than its current value.
The application of Floyd's algorithm to the graph in Figure 8.5 is illustrated
in Figure 8.7.
Here is a pseudocode of Floyd's algorithm. It takes advantage of the fact that
the next matrix in sequence (8.8) can be written over its predecessor.
Floyd(W[l..n, l..n])
//Implements Floyd's algorithm for the all-pairs shortest -paths problem
//Input: The weight matrix W of a graph with no negative-length cycle
//Output: The distance matrix of the shortest paths' lengths
D <-- W /lis not necessary if W can be overwritten
for k <-- 1 to n do
for i <-- 1 to n do
for j <-- 1 to n do
D[i. j] <-- min{D[i, j], D[i, k] + D[k, j])
retnrn D
ALGORITHM
Obviously, the time efficiency of Floyd's algorithm is cubic-as is the time
efficiency of Warshall's algorithm. In the next chapter, we examine Dijkstra's
algorithm-another method for finding shortest paths.
8.2
Warshall's and Floyd's Algorithms
~ T ~~
a
b
c
d
3
7
OIOl~
1
b
2
c
=
d
6
a
0
7
0
b
r~
=
=
c
0
7
3
5
0
9
II
a
c
d
b
Ol1l~
2
d
6
c
~
b
T Il
r l~
T 1]
b
c
d
2
9
6
a
0131~
b
c
d
2
9
6
a
0141~
b
2
c
7
d
6
0
7
5
0
=~
b c
10 3
0 5
7 0
16 9
d
c
d
b
10
0
7
16
3
5
0
9
Lengths of the shortest paths
with no intermediate vertices
(O(Ol is simply the weight matrix).
d
='3
012}~
291
Lengths of the shortest paths
with intermediate vertices numbered
not higher than 1, i.e. just a
(note two new shortest paths from
b to c and from d to c).
Lengths of the shortest paths
with intermediate vertices numbered
not higher than 2, i.e. a and b
(note a new shortest path from c to a).
Lengths of the shortest paths
with intermediate vertices numbered
not higher than 3, i.e. a, b, and c
(note four new shortest paths from a to b,
from a to d, from b to d, and from d to b).
Lengths of the shortest paths
with intermediate vertices numbered
not higher than 4, i.e. a, b, c, and d
{note a new shortest path from c to a).
FIGURE 8.7 Application of Floyd's algorithm to the graph shown. Updated elements are
shown in bold.
We finish this section with an important general comment. It deals with a general principle that underlines dynamic programming algorithms for optinlization
problems. Richard Bellman called it the principle of optimality. In terms somewhat different from its original formulation, it says that an optimal solution to
any instance of an optimization problem is composed of optimal solutions to its
subinstances. The principle of optimality holds more often than not. (To give a
rather rare example, it fails for finding longest simple paths.) Although its applicability to a particular problem needs to be verified, of course, such a verification
is usually not a principal difficulty in developing a dynamic programming algorithm. The challenge typically lies in figuring out what smaller subinstances need
to be considered and in deriving an equation relating a solution to any instance
to solutions to its smaller subinstances. We consider a few more examples in the
remaining sections of this chapter and their exercises.
:
f
292
Dynamic Programming
-----Exercises 8 . 2 - - - - - - - - - - - - - - - 1. Apply Warshall"s algorithm to find the transitive closure of the digraph defined by the following adjacency matrix
2. a. Prove that the time efficiency of Warshall's algorithm is cubic.
b. Explain why the time efficiency of Warshall's algorithm is inferior to that
of the traversal-based algorithm for sparse graphs represented by their
adjacency lists.
Explain how to implement Warshall's algorithm without using extra memory
for storing elements of the algorithm's intermediate matrices.
4. Explain how to restructure the innermost loop of the algorithm Warshall to
make it run faster at least on some inputs.
5. Rewrite the pseudocode of Warshall's algorithm assuming that the matrix
rows are represented by bit strings on which the bitwise or operation can be
performed.
6. a. Explain how Warshall's algorithm can be used to determine whether a
given digraph is a dag (directed acyclic graph). Is it a good algorithm for
this problem?
b. Is it a good idea to apply Warshall's algorithm to find the transitive closure
of an undirected graph?
7. Solve the all-pairs shortest-path problem for the digraph with the weight
matrix
~ ~
00
00
00
00
00
3
0
2
3
00
00
[
1
2
4
0
00
7]
8. Prove that the next matrix in sequence (8.8) of Floyd's algorithm can be
written over its predecessor.
{'!.\ Give an example of a graph or a digraph with negative weights for which
Floyd's algorithm does not yield the correct resnlt.
._1«: Enhance Floyd's algorithm so that shortest paths themselves, not just their
·
lengths, can be found.
8.3
Optimal Binary Search Trees
293
11. Jack Straws In the game of Jack Straws, a number of plastic or wooden
"straws" are dumped on the table and players try to remove them one-byone without disturbing the other straws. Here, we are only concerned with
whether various pairs of straws are connected by a path of touching straws.
Given a list of the endpoints for n > 1 straws (as if they were dumped on a large
piece of graph paper), determine all the pairs of straws that are connected.
Note that touching is connecting, but also that two straws can be connected
indirectly via other connected straws [1994 East-Central Regionals of the
ACM International Collegiate Programming Contest].
8.3 Optimal Binary Search Trees
A binary search tree is one of the most important data structures in computer
science. One of its principal applications is to implement a dictionary, a set of
elements with the operations of searching, insertion, and deletion. If probabilities
of searclting for elements of a set are known (e.g., from accumulated data about
past searches), it is natural to pose a question about an optimal binary search
tree for which the average number of comparisons in a search is the smallest
possible. (For simplicity, we limit our discussion to minimizing the average number
of comparisons in a successful search. The method can be extended to include
unsuccessful searches as well.)
As an example, consider four keys A, B, C, and D to be searched for with
probabilities 0.1, 0.2, 0.4, and 0.3, respectively. Figure 8.8 depicts two out of
14 possible binary search trees containing these keys. The average number of
comparisons in a successful search in the first of these trees is 0.1·1 + 0.2·2 + 0.4-3
+ 0.3-4 = 2.9, while for the second one it is 0.1·2 + 0.2·1 + 0.4-2 + 0.3-3 = 2.1. Neither
of these two trees is, in fact, optimal. (Can you tell which binary tree is optimal?)
A
8
A
8
c
c
0
0
FIGURE 8.8 Two out of 14 possible binary search trees with keys A, B, C,
and D
294
Dynamic Programming
Optimal
BST for
Optimal
BST for
a;, ... , ak-1
ak+ 1 , ... , ai
FIGURE 8.9 Binary search tree (BSTI with root
Ti
k-1
an
d Tj
a, and two optimal binary search subtrees
k+1
For our tiny example, we could find the optimal tree by generating all 14
binary search trees with these keys. As a general algorithm, this exhaustive-search
approach is unrealistic: the total number of binary search trees with n keys is equal
to the nth Catalan number
2
1
n+1
c(n) = ( n) -
n
for n > 0,
c(O) = 1,
which grows to infinity as fast as 4" I nl. 5 (see Problem 7 in the exercises).
So let a 1•... , a,. be distinct keys ordered from the smallest to the largest and
let p 1 , ... , p,. be the probabilities of searching for them. Let C[i, j] be the smallest
average number of comparisons made in a successful search in a binary search tree
made up of keys ai, ... , aj, where i, j are some integer indices, 1::; i:::; j::::; n.
Thus, following the classic dynamic programming approach, we will find values of
C[i, j] for all smaller instances of the problem, although we are interested just in
C[1, n ]. To derive a recurrence underlying the dynamic programming algorithm,
we will consider all possible ways to choose a root ak among the keys a,, ... , aj.
For such a binary search tree (Figure 8.9), the root contains key ak> the left subtree
r,'- 1 contains keys a;, ... , 1 optimally arranged, and the right subtree T/+1
contains keys ak+ 1, ...• aj also optimally arranged. (Note how we are taking
advantage of the principle of optimality here.)
If we count tree levels starting with 1 (to make the comparison numbers equal
the keys' levels), the following recurrence relation is obtained:
T/
a,_
8.3
295
Optimal Binary Search Trees
k-1
L p, ·(level of a, in rk-l + 1)
C(i, j] = min{pk · 1 +
I;Sk:SJ
.
I
S=l
j
L p_, ·(level of a, in T./+ + 1)}
+
1
s=k+l
k-1
k-1
1 + '\' P.,
= min{pk + '\' p,. ·level of a, in rk1
0
~
t:SkS)
~·=l
S=/
j
j
L p,·levelofa,inT.{H+ L p,)
+
s=k+1
.v=k+l
k-1
j
-
j
L Ps ·level of a in Tk~l + L Ps}
= i~i~.{L P.1• ·level of as in T/- +
1
5
s=k+ 1
_J s=i
s=i
j
= min{C[i, k -1]+ C[k + 1, J]) +
t;Sk;S:j
L P.,.
.
S=l
Thus, we have the recurrence
}
C(i, j] = min(C[i, k -1] + C[k + 1, j]) +
t:Sk:Sj
L p, for 1:5; :5 j :5 n. (8.11)
.
S=l
We assume in formula (8.11) that C[i, i- 1] = 0 for 1 :5 i :5 n + 1, which cau be
interpreted as the number of comparisons in the empty tree. Note that this formula
implies that
C[i, i] =Pi
for 1 :5 i :5 n,
as it should be for a one-node binary search tree containing a;.
The two-dimensional table in Figure 8.10 shows the values needed for computing C(i, j] by formula (8.11): they are in row i and the columns to the left of
column j and in column j and the rows below row i. The arrows point to the
pairs of entries whose sums are computed in order to find the smallest one to be
recorded as the value of C[i, j]. This suggests filling the table along its diagonals,
starting with all zeros on the main diagonal and given probabilities Pi, 1 :5 i :5 n,
right above it and moving toward the upper right corner.
The algorithm we sketched computes C[1, n ]-the average number of comparisons for successful searches in the optimal binary tree. If we also want to get the
optimal tree itself, we need to maintain another two-dimensional table to record
the value of k for which the minimum in (8.11) is achieved. The table has the same
shape as the table in Figure 8.10 and is filled in the same manner, starting with
entries R[i, i] = i for 1 :5 i :5 n. When the table is filled, its entries indicate indices
of the roots of the optimal subtrees, which makes it possible to reconstruct an
optimal tree for the entire set given.
296
Dynamic Programming
n
j
0
goal
P,
0
p2
0
Cli,il
t
\_
'--
I-+
Pn
n +1
0
FIGURE 8.10 Table of the dynamic programming algorithm for constructing an optimal
binary search tree
EXAMPLE Let us illustrate the algorithm by applying it to the four-key set we
used at the beginning of this section:
key
probability
A
B
c
D
0.1
0.2
0.4
0.3
The initial tables look like this:
0
0
2
3
main table
3
2
0.1
0
2
0.2
0
4
5
0
4
3
0.4
0
0.3
0
4
root table
1 2 3
4
2
3
4
5
Let us compute C[1, 2]:
. k = 1: C[1, o] + C[2, 21 + L:;d p, = o + 0.2 + o.3 = o.s
C[1, 2] = mm
= 0.4.
2
k = 2: C[1, 1] + C[3, 2] + L:,~l p, = 0.1 + 0 + 0.3 = 0.4
8.3
Optimal Binary Search Trees
297
c
B
D
A
FIGURE 8.11 Optimal binary search tree for the example
Thus, out of two possible binary trees containing the first two keys, A and B, the
root of the optimal tree has index 2 (i.e., it contains B), and the average number
of comparisons in a successful search in this tree is 0.4.
We will ask you to finish the computations in the exercises. You should arrive
at the following final tables:
0
0
2
3
4
5
main table
2
3
0.1 0.4 1.1
0
0.2 0.8
0
0.4
0
root table
4
1.7
1.4
1.0
0.3
0
0
2
3
4
5
1
2
2
2
3
3
3
3
4
3
3
3
4
Thus, the average number of key comparisons in the optimal tree is equal to
1.7. Since R[1, 4] = 3, the root of the optimal tree contains the third key, i.e., C. Its
left subtree is made up of keys A and B, and its right subtree contains just key D
(why?). To find the specific structure of these subtrees, we find first their roots by
consulting the root table again as follows. Since R[l, 2] = 2, the root ofthe optimal
tree containing A and B is B, with A being its left child (and the root of the onenode tree: R[1, 1] = 1). Since R[4, 4] = 4, the root of this one-node optimal tree is
its only key D. Figure 8.11 presents the optimal tree in its entirety.
113
Here is a pseudocode of the dynamic programming algorithm.
ALGORITHM
OptimalBST(P[l..n])
//Finds an optimal binary search tree by dynamic programming
//Input: An array P[l..n] of search probabilities for a sorted list of n keys
//Output: Average number of comparisons in successful searches in the
II
optimal EST and table R of subtrees' roots in the optimal EST
for i +- 1 to n do
C[i, i -1]+-0
C[i, i] +- P[i]
R[i, i] +- i
298
Dynamic Programming
C[n + 1, n] <- 0
for d <- 1 to n - 1 do //diagonal count
for i <- 1 to n - d do
}<-i+d
minval +- oo
for k <- i to j do
ifC[i, k -1] + C[k + 1, j] < minval
min val<- C[i, k- 1] + C[k + 1, }]; kmin <- k
R[i, j] <- kmin
sum <- P[i]; for s <- i + 1 to j do sum <-sum+ P[s]
C[i, j] <-min val+ sum
return C[1, n ], R
The algoritlun's space efficiency is clearly quadratic; the time efficiency of this
version of the algoritlun is cubic (why?). A more careful analysis shows that entries
in the root table are always nondecreasing along each row and column. This limits
values for R[i, j] to the range R[i, j - 1], ... , R[i + 1, j] and makes it possible to
reduce the running time of the algorithm to 8(n 2 ).
-----Exercises 8 . 3 - - - - - - - - - - - - - - - 1. Finish the computations started in the section's example of constructing an
optimal binary search tree.
2. a. Why is the time efficiency of algorithm Optima/EST cubic?
b. Why is the space efficiency of algoritlun Optima/EST quadratic?
3. Write a pseudocode for a linear-time algorithm that generates the optimal
binary search tree from the root table.
4. Devise a way to compute the smns I:~~; p,, which are used in the dynantic
programming algoritlun for constructing an optimal binary search tree, in
constant time (per sum).
5. True or false: The root of an optimal binary search tree always contains the
key with the highest search probability?
6. How would you construct an optimal binary search tree for a set of n keys if
all the keys are equally likely to be searched for? What will be the average
number of comparisons in a successful search in such a tree if n = 2k?
7. a. Show that the number of distinct binary search trees b(n) that can be
constructed for a set of n orderable keys satisfies the recurrence relation
n-1
b(n) =
L b(k)b(n- 1- k) for n > 0, b(O) = 1.
k~O
8.4
The Knapsack Problem and Memory Functions
299
b. It is known that the solution to this recurrence is given by the Catalan
numbers. Verify this assertion for n = 1, 2, ... , 5.
c. Find the order of growth of b(n). What implication does the answer to
this question have for the exhaustive-search algorithm for constructing an
optimal binary search tree?
8. Design a e (n 2 ) algorithm for finding an optimal binary search tree.
9. Generalize the optimal binary search algorithm by taking into account unsuccessful searches.
~~ Matrix chain multiplication Consider the problem of minimizing the total
- number of multiplications made in computing the product of n matrices
At· A2 ..... An
whose dimensions are do by dt, dt by d2, ... , dn-t by dn, respectively. Assume
that all intermediate products oftwo matrices are computed by the brute-force
(definition-based) algorithm.
a. Give an example of three matrices for which the number of multiplications
in (At· A 2) · A 3 and A 1 · (A 2 · A 3) differ at least by a factor 1000.
b. How many different ways are there to compute the chained product of n
matrices?
c. Design a dynamic programming algorithm for finding an optimal order of
multiplying n matrices.
8.4 The Knapsack Problem and Memory Functions
We start this section with designing the dynamic programming algorithm for
the knapsack problem: given n items of known weights Wt, ... , wn and values
Vt, ... , vn and a knapsack of capacity W, find the most valuable subset of the items
that fit into the knapsack. (Tiris problem was introduced in Section 3.4, where we
discussed solving it by an exhaustive-search algorithm.) We assume here that all
the weights and the knapsack's capacity are positive integers; the item values do
not have to be integers.
To design a dynamic programming algorithm, we need to derive a recurrence
relation that expresses a solution to an instance of the knapsack problem in terms
of solutions to its smaller subinstances. Let us consider an instance defined by the
first i items, 1:::; i.::::; n, with weights w 1, ... , wi, values v1, ... , vi, and knapsack
capacity j, 1 S j S W. Let V[i, jJ be the value of an optimal solution to this
instance, i.e., the value of the most valuable subset of the first i items that fit into
the knapsack of capacity j. We can divide all the subsets of the first i items that fit
the knapsack of capacity j into two categories: those that do not include the ith
item and those that do. Note the following:
;:
300
Dynamic Programming
0
0
j-wi
0
j
w
0
0
i-1
0
0
Vli-1, j-w;l
Vli-1, il
n
0
0
wi, vi
VIi, il
goal
FIGURE 8.12 Table for solving the knapsack problem by dynamic programming
1.
2.
Among the subsets that do not include the ith item, the value of an optimal
subset is, by definition, V[i - 1, j].
Among the subsets that do include the ith item (hence, j - w; ?: 0), an optimal
subset is made up of this item and an optimal subset of the first i - 1 items that
fit into the knapsack of capacity j - w;. The value of such an optimal subset
is v; + V[i - 1, j - w;].
Thus, the value of an optimal solution among all feasible subsets of the first i
items is the maximum of these two values. Of course, if the ith item does not fit
into the knapsack, the value of an optimal subset selected from the first i items
is the same as the value of an optimal subset selected from the first i - 1 items.
These observations lead to the following recurrence:
[
.
v l, Jl =
{ max{ V[i - 1, j], V; + V[i - 1, j. -. w. ;])
V[i -1, J]
'
l
if j - w; ?: 0
if j - W; < 0. (S.U)
It is convenient to define the initial conditions as follows:
V[O, j] = 0 for j?: 0 and V[i, 0] = 0 fori?: 0.
(8.13)
Our goal is to find V[n, W], the maximal value of a subset of then given items that
fit into the knapsack of capacity W, and an optimal subset itself.
Figure 8.12 illustrates the values involved in equations (8.12) and (8.13). For
i, j > 0, to compute the entry in the ith row and the jth column, V[i, J], we
compute the maximum of the entry in the previous row and the same column
and the sum of v; and the entry in the previous row and w; columns to the left.
The table can be filled either row by row or column by column.
EXAMPLE 1 Let us consider the instance given by the following data:
item
weight
value
1
2
3
4
2
1
3
2
$12
$10
$20
$15
capacity W = 5
8.4 The Knapsack Problem and Memory Functions
301
capacity j
0
w1 = 2, v1 = 12
w2 = 1, v2 = 10
w3 ~ 3, v3 ~ 20
2
3
W4=2,V4=15
4
0
0
0
0
0
0
0
0
10
10
10
2
0
12
12
12
15
3
0
12
22
22
25
4
0
12
22
30
30
5
0
12
22
32
37
FIGURE 8.13 Example of solving an instance of the knapsack problem by the dynamic
programming algorithm
The dynamic programming table, filled by applying formulas (8.12) and (8.13), is
shown in Figure 8.13.
Thus, the maximal value is V[4, 5] = $37. We can find the composition of an
optimal subset by tracing back the computations of this entry in the table. Since
V[4, 5] I= V[3, 5], item 4 was included in an optimal solution along with an optimal
subset for filling 5 - 2 = 3 remaining units of the knapsack capacity. The latter is
represented by element V[3, 3]. Since V[3, 3] = V[2, 3], item 3 is not a part of an
optimal subset. Since V[2, 3] I= V[l, 3], item 2 is a part of an optimal selection,
which leaves element V[l, 3 -1] to specify its remaining composition. Similarly,
since V[1, 2] 1= l?[O, 2], item 1 is the final part of the optimal solution {item 1, item
2, item 4).
Ill
The time efficiency and space efficiency of this algorithm are both in 8 (n W).
The time needed to find the composition of an optimal solution is in O(n + W).
You are asked to prove these· assertions in the exercises.
Memory Functions
As we discussed at the beginning of this chapter and illustrated in subsequent
sections, dynamic programming deals with problems whose solutions satisfy a
recurrence relation with overlapping subproblems. The direct top-down approach
to finding a solution to such a recurrence leads to an algorithm that solves common
subproblems more than once and hence is very inefficient (typically, exponential
or worse). The classic dynamic programming approach, on the other hand, works
bottom-up: it fills a table with solutions to all smaller subproblems, but each of
them is solved only once. An unsatisfying aspect of this approach is that solutions
to some of these smaller subproblems are often not necessary for getting a solution
to the problem given. Since this drawback is not present in the top-down approach,
it is natural to try to combine the strengths of the top-down and bottom-up
approaches. ~.ggf!l!~.t 0 geta.method thatsolves.only subproblems.which.'!fe.
':'ec!'S§'lJ)randdoesitonlyonce.-Such a method exists; it is based on using memory
fUnctions [Bra96].
~m'
~.·: ·. 1{·: 1
itlili
,, '!!
;.'1!
1.'.!.·. ilii
302
Dynamic Programming
This method solves a given problem in the top-down manner but, in addition,
maintains a table of the kind that would have been used by a bottom-up dynamic
programming algorithm. Initially, all the table's entries are initialized with a special "null" symbol to indicate that they have not yet been calculated. Thereafter,
whenever a new value needs to be calculated, the method checks the corresponding entry in the table first: if this entry is not "null," it is simply retrieved from the
table; otherwise, it is computed by the recursive call whose result is then recorded
in the table.
The following algorithm implements this idea for the knapsack problem. After
initializing the table, the recursive function needs to be called with i = n (the
number of items) and j = W (the capacity of the knapsack).
ALGORITHM
MFKnapsack(i, j)
//Implements the memory function method for the knapsack problem
//Input: A nonnegative integer i indicating the number of the first
II
items being considered and a nonnegative integer j indicating
II
the knapsack's capacity
//Output: The value of an optimal feasible subset of the first i items
//Note: Uses as global variables input arrays Weights[l .. n ], V alues[l .. n ],
//and table V[O .. n, O.. W] whose entries are initialized with -l's except for
//row 0 and column 0 initialized with O's
ifV[i,j]<O
if j < Weights[i]
value<-- MFKnapsack(i- 1, j)
else
value<-- max(MFKnapsack(i- 1, j),
Values[i] + MFKnapsack(i - 1, j - Weights[i]l)
V[i, j] <--value
return V[i, j]
EXAMPLE 2 Let us apply the memory function method to the instance considered in Example 1. Figure 8.14 gives the results. Only 11 out of20 nontrivial values
(i.e., not those in row 0 or in column 0) have been computed. Just one nontrivial
entry, V[1, 2], is retrieved rather than being recomputed. For larger instances, the
proportion of such entries can be significantly larger.
II!
In general, we cannot expect more than a constant-factor gain in using the
memory function method for the knapsack problem because its time efficiency
class is the same as that of the bottom-up algorithm (why?). A more significant
improvement can be expected for dynamic programming algorithms in which a
computation of one value takes more than constant time. You should also keep
in mind that a memory function method may be less space-efficient than a spaceefficient version of a bottom-up algorithm.
8.4 The Knapsack Problem and Memory Functions
303
capacity j
i
0
w1 = 2, v1 = 12
W2= 1, V2= 10
W3
= 3, v3 = 20
W4 = 2, v4 =
15
2
3
4
0
0
0
0
0
0
0
0
2
0
12
12
3
0
12
22
22
4
0
12
5
0
12
22
32
37
FIGURE 8.14 Example of solving an instance of the knapsack problem by the memory
function algorithm
-----Exercises 8 . 4 - - - - - - - - - - - - - - 1. a. Apply the bottom-up dynamic programming algorithm to the following
instance of the knapsack problem:
item
weight
value
1
2
3
3
2
1
$25
$20
$15
4
4
$40
5
5
$50
capacity W = 6.
-<I
b. How many different optimal subsets does the instance of part (a) have?
c. In general, how can we use the table generated by the dynamic programming algorithm to tell whether there is more than one optimal subset for
the knapsack problem's instance?
,•. \ .'· · /·
2. a. Write a pseudocode of the bottom-up dynamic programming algorithm for
the knapsack problem.
b. Write a pseudocode of the algorithm that finds the composition of an
optimal subset from the table generated by the bottom-up dynamic pro:· gramming algorithm for the knapsack problem.
3. For the bottom-up dynamic programming algorithm for the knapsack problem, prove that
a. its time efficiency is in El(nW).
b. its space efficiency is in El(nW).
c. the time needed to find the composition of an optimal subset from a filled
dynamic programming table is in O(n + W).
4. a. True or false: A sequence of values in a row of the dynamic programming
table for an instance of the knapsack problem is always nondecreasing?
304
Dynamic Programming
b. True or false: A sequence of values in a column of the dynamic programming table for an instance of the knapsack problem is always nondecreasing?
5. Apply the memory function method to the instance of the knapsack problem
given in Problem 1. Indicate the entries of the dynamic programming table
that are: (i) never computed by the memory function method on this instance;
(ii) retrieved without a recomputation.
6. Prove that the efficiency class of the memory function algorithm for the knapsack problem is the same as tbat of the bottom-up algorithm (see Problem 3).
~rite a pseudocode of a memory function for the optimal binary search tree
Vrroblem. (You may limit your function to finding the smallest number of key
comparisons in a successful search.)
8. Give two reasons why the memory function approach is unattractive for the
problem of computing a binomial coefficient.
9. Design a dynamic programming algorithm for the change-making problem:
given an amount n and unlimited quantities of coins of each of the denominations d 1, d2 , ... , dm, find the smallest number of coins that add upton or
indicate that the problem does not have a solution.
10. Write a research report on one of the following well-known applications of
dynamic programming:
a. finding the longest common subsequence in two sequences
b. optimal string editing
c. minimal triangulation of a polygon
SUMMARY
"
Dynamic programming is a technique for solving problems with overlapping
subproblems. Typically, these subproblems arise from a recurrence relating
a solution to a given problem with solutions to its smaller subproblems of the
same type. Dynamic programming suggests solving each smaller subproblem
once and recording the results in a table from which a solution to the original
problem can be then obtained.
"
Applicability of dynamic programming to an optimization problem requires
the problem to satisfy the principle of optimality: an optimal solution to any
of its instances must be made up of optimal solutions to its subinstances.
"'
Computing a binomial coefficient via constructing the Pascal triangle can
be viewed as an application of the dynamic programming technique to a
nonoptirnization problem.
Summary
305
,.
Warshall's algorithm for finding the transitive closure and Floyd's algorithm
for the all-pairs shortest -paths problem are based on the idea that can be
interpreted as an application of the dynamic programming technique.
,.
Dynamic programming can be used for constructing an optimal binary search
tree for a given set of keys and known probabilities of searching for them.
"'
Solving a knapsack problem by a dynamic programming algorithm exemplifies an application of this technique to difficult problems of combinatorial
optimization.
,.
The memory function technique seeks to combine strengths of the topdown and bottom-up approaches to solving problems with overlapping
subproblems. It does this by solving, in the top-down fashion but only once,
just necessary subproblems of a given problem and recording their solutions
in a table.
l'/1
Greedy Technique
Greed, for lack of a better word, is good! Greed is right! Greed works!
-Michael Douglas, U.S. actor in the role of Gordon Gecko,
in the film Wall Street, 1987
L
et us start with the change-making problem faced by millions of cashiers all
over the world (at least subconsciously): give change for a specific amount
n with the least number of coins of the denominations d 1 > d 2 > ... > dm used
in that locale. For example, the widely used coin denominations in the United
States are d1 = 25 (quarter), d2 = 10 (dime), d3 = 5 (nickel), and d4 = 1 (penny).
How would you give change with coins of these denominations of, say, 48 cents?
If you came up with the answer 1 quarter, 2 dimes, and 3 pennies, yon followedconsciously or not-a logical strategy of making a sequence of best choices among
the currently available alternatives. Indeed, in the first step, you could have given
one coin of any of the four denominations. "Greedy" thinking leads to giving one
quarter because it reduces the remaining amount the most, namely, to 23 cents. In
the second step, you had the same coins at your disposal, but you could not give
a quarter because it would have violated the problem's constraints. So your best
selection in this step was one dime, reducing the remaining amount to 13 cents.
Giving one more dime left you with 3 cents to be given with three pennies.
Is this solution to the instance of the change-making problem optimal? Yes, it
is. In fact, it is possible to prove that the greedy algorithm yields an optimal solution
for every positive integer amount with these coin denominations. At the same
time, it is easy to give an example of "weird" coin denominations-e.g., d1 = 7,
d2 = 5, d3 = 1-that may not yield an optimal solution for some amounts. (It is
the reason you were asked to develop a dynamic programming algorithm for this
problem in Exercises 8.4: that algorithm works for any set of coin denominations
by either returning an optimal solution or reporting that no solution exists.)
The approach applied in the opening paragraph to the change-making problem is called greedy. Computer scientists consider it a general design technique
despite the fact that it is applicable to optimization problems only. The greedy
307
__,.
I
Greedy Technique
308
approach suggests constructing a solution through a sequence of steps, each expanding a partially constructed solution obtained so far, until a complete solution
to the problem is reached. On each step-and this is the central point of this
technique-the choice made must be
''""'
.,,,,..
, '"'"''"'"'
feasible, i.e., it has to satisfy the problem's constraints
locally optimal, i.e., it has to be the best local choice among all feasible choices
available on that step
irrevocable, i.e., once made, it cannot be changed on subsequent steps of the
algorithm
These requirements explain the technique's name: on each step, it suggests
a "greedy" grab of the best alternative available in the hope that a sequence
of locally optimal choices will yield a (globally) optimal solution to the entire
problem. We refrain from a philosophical discussion of whether greed is good
or bad. (If you have not seen the movie from which the chapter's epigraph is
taken, its hero did not end up well.) From our algorithmic perspective, the question is whether a greedy strategy works or not. As we shall see, there are problems for which a sequence of locally optimal choices does yield an optimal solution for every instance of the problem in question. However, there are others
for which this is not the case; for such problems, a greedy algorithm can still
be of value if we are interested in or have to be satisfied with an approximate
solution.
In the first two sections of the chapter, we discuss two classic algorithms for the
minimum spanning tree problem: Prim's algorithm and Kruskal's algorithm. What
is remarkable about these algorithms is the fact that they solve the same problem
by applying the greedy approach in two different ways, and both of them always
yield an optimal solution. In Section 9.3, we introduce another classic algorithmDijkstra's algorithm for the shortest-path problem in a weighted graph. Section 9.4
is devoted to Huffman trees and their principal application, Huffman codes-an
important data compression method that can be interpreted as an application of
the greedy technique. Finally, a few examples of approximation algorithms based
on the greedy approach are discussed in Section 12.3.
As a rule, greedy algorithms are both intuitively appealing and simple. Despite
their apparent simplicity, there is a rather sophisticated theory behind the technique, which is based on the abstract combinatorial structure called "matroid."
We are not going to discuss it here; an interested reader can check such sources
as [Cor01].
9.1
Prim's Algorithm
The following problem arises naturally in several practical situations: given n
points, connect them in the cheapest possible way so that there will be a path
between every pair of points. We can represent the points by vertices of a graph,
9.1
Prim's Algorithm
309
possible connections by the graph's edges, and the connection costs by the edge
weights. Then the question can be posed as the minimum spanning tree problem,
defined formally as follows.
DEFINITION A spanning tree of a connected graph is its connected acyclic
subgraph (i.e., a tree) that contains all the vertices of the graph. A minimum
spanning tree of a weighted connected graph is its spanning tree of the smallest
weight, where the weight of a tree is defined as the sum of the weights on all its
edges. The minimum spanning tree problem is the problem of finding a minimum
spanning tree for a given weighted connected graph.
Figure 9.1 presents a simple example illustrating these notions.
If we were to try an exhaustive-search approach to constructing a minimum
spanning tree, we would face two serious obstacles. First, the number of spanning
trees grows exponentially with the graph size (at least for dense graphs). Second,
generating all spanning trees for a given graph is not easy; in fact, it is more difficult
than finding a minimum spanning tree for a weighted graph by using one of several
efficient algorithms available for this problem. In this section, we outline Prim's
algorithm, which goes back to at least 1957 [Pri57].
Prim's algorithm constructs a minimum spanning tree through a sequence
of expanding subtrees. The initial subtree in such a sequence consists of a single
vertex selected arbitrarily from the set V of the graph's vertices. On each iteration,
we expand the current tree in the greedy manner by simply attaching to it the
nearest vertex not in that tree. (By the nearest vertex, we mean a vertex not in
the tree connected to a vertex in the tree by an edge of the smallest weight.~
can be broken arbitrarily.) The algorithm stops after all the graph's vertices have
been included in the tree being constructed. Since the algorithm expands a tree
by exactly one vertex on each of its iterations, the total number of such iterations
is n - 1, where n is the number of vertices in the graph. The tree generated by the
algorithm is obtained as the set of edges used for the tree expansions.
~
graph
r:
w(T2 ) ~ 9
w(T3 ) ~ 8
FIGURE 9.1 Graph and its spanning trees; T1 is the minimum spanning tree
Greedy Technique
310
Here is a pseudocode of this algorithm.
ALGORITHM Prim(G)
//Prim's algorithm for constructing a minimum spanning tree
//Input: A weighted connected graph G = (V, E)
//Output: Er, the set of edges composing a minimum spanning tree of G
Vr <- {v0 ]
//the set of tree vertices can be initialized with any vertex
Er <- 0
for i <- 1 to IV I - 1 do
find a minimum-weight edge e* = (v*, u') among all the edges (v, u)
such that vis in Vr and u is in V - Vr
Vr <- Vr U {u*]
Er <- Er U {e*]
return Ey
The nature of Prim's algorithm makes it necessary to provide each vertex not
in the current tree with the information about the shortest edge connecting the
vertex to a tree vertex. We can provide such information by attaching two labels
to a vertex: the name of the nearest tree vertex and the length (the weight) of the
corresponding edge. Vertices that are not adjacent to any of the tree vertices can
be given the oo label indicating their "infinite" distance to the tree vertices and
a null label for the name of the nearest tree vertex. (Alternatively, we can split
the vertices that are not in the tree into two sets, the "fringe" and the "unseen."
The fringe contains only the vertices that are not in the tree but are adjacent to at
least one tree vertex. These are the candidates from which the next tree vertex
is selected. The unseen vertices are all the other vertices of the graph, called
"unseen" because they are yet to be affected by the algorithm.) With such labels,
finding the next vertex to be added to the current tree T = (VT, Er) becomes a
simple task of finding a vertex with the smallest distance label in the set V- Vy.
Ties can be broken arbitrarily.
After we have identified a vertex u* to be added to the tree, we need to perform
two operations:
Ill
m
Move u* from the set V - Vr to the set of tree vertices Vr.
For each remaining vertex u in V - v,. that is connected to u* by a shorter
edge than the u's current distance label, update its labels by u* and the weight
of the edge between u* and u, respectively1
Figure 9.2 demonstrates the application of Prim's algorithm to a specific graph.
Does Prim's algorithm always yield a minimum spanning tree? The answer
to this question is yes. Let us prove by induction that each of the subtrees T,,
i = 0, ... , n- 1, generated by Prim's algorithm is a part (i.e., a subgraph) of some
1.
If the implementation with the fringe-unseen split is pursued, all the unseen vertices adjacent to u*
must also be moved to the fringe.
;c,,p;
'
3
Tree vertices
4
Remaining vertices
a(-, -)
b(a, 3) c(-, oo) d(-, oo)
e(a, 6) f(a, 5)
b(a, 3)
c(b, 1) d(-, oo) e(a, 6)
f(b, 4)
c(b, 1)
d(c, 6) e(a, 6) f(b, 4)
f(b, 4)
d(f, 5) e(f, 2)
e(f, 2)
d(f, 5)
4
6
Illustration
d(f, 5)
FIGURE 9.2 Application of Prim's algorithm. The parenthesized labels of a vertex in the
middle column indicate the nearest tree vertex and edge weight; selected
vertices and edges are shown in bold.
311
312
Greedy Technique
v'
v
---------e,
u'
u
FIGURE 9.3 Correctness proof of Prim's algorithm
minimum spanning tree. (This immediately implies, of course, that the last tree in
the sequence, T11 _ 1, is a minimum spanning tree itself because it contains all n
vertices of the graph.) The basis of the induction is trivial, since T0 consists of a
single vertex and hence must be a part of any minimum spanning tree. For the
inductive step, let us assume that Ti~1 is part of some minimum spanning tree T.
We need to prove that T;, generated from T; _1 by Prim's algorithm, is also a part
of a minimum spanning tree. We prove this by contradiction by assuming that no
minimum spanning tree of the graph can contain T1• Let e1 = (v, u) be the minimum
weight edge from a vertex in T; _ 1 to a vertex not in T1_ 1 used by Prim's algorithm to
expand T1_ 1 to 1;. By our assumption, e1 cannot belong to the minimum spanning
tree T. Therefore, if we add e; to T, a cycle must be formed (Figure 9.3).
In addition to edge e1 = (v, u), this cycle must contain another edge (v', u')
connecting a vertex v' E T1_ 1 to a vertex u' that is not in T1_ 1. (It is possible that
v' coincides with v or u' coincides with u but not both.) If we now delete the edge
(v', u') from this cycle, we obtain another spanning tree of the entire graph whose
weight is less than or equal to the weight ofT since the weight of e1 is less than or
equal to the weight of (v', u'). Hence, this spanning tree is a minimum spanning
tree, which contradicts the assumption that no minimum spanning tree contains
T1• This completes the correctness proof of Prim's algorithm.
How efficient is Prim's algorithm? The answer depends on the data structures
chosen for the graph itself and for the priority queue of the set V - Vr whose
vertex priorities are the distances to the nearest tree vertices. (You may want
to take another look at the example in Figure 9.2 to see that the set V - Vr
indeed operates as a priority queue.) In particular, if a graph is represented by
its weight matrix and the priority queue is implemented as an unordered array,
the algorithm's running time will be in El(IV[ 2). Indeed, on each of the [V[-1
iterations, the array implementing the priority queue is traversed to find and delete
the minimum and then to update, if necessary, the priorities of the remaining
vertices.
We can also implement the priority queue as a min-heap. A min-heap is a
mirror image of the heap structure discussed in Section 6.4. (In fact, it can be implemented by constructing a heap after negating all the key values given.) Namely,
a min-heap is a complete binary tree in which every element is less than or equal
~
9.1
I
Prim's Algorithm
313
to its children. All the principal properties of heaps remain valid for min-heaps,
with some obvious modifications. For example, the root of a min-heap contains the
smallest rather tban the largest element. Deletion of the smallest element from
and insertion of a new element into a min-heap of size n are O(log n) operations,
and so is the operation of changing an element's priority (see Problem 10).
If a graph is represented by its adjacency lists and the priority queue is implemented as a min-heap, the running time of the algorithm is in 0(1 E I log 1VI). 1bis
is because the algorithm performs lVI- 1 deletions of the smallest element and
makes lEI verifications and, possibly, changes of an element's priority in a minheap of size not greater than IV 1. Each of these operations, as noted earlier, is a
O(log lVI) operation. Hence, the running time of this implementation of Prim's
algorithm is in
(IVI-1 + IEIJO(log lVI) = O(!Eilog JVI)
because, in a connected graph, lVI -1 :S lEI.
In the next section, you will find another greedy algorithm for the minimum
spanning tree problem, which is "greedy" in a manner different from that of Prim's
algorithm.
-----Exercises 9.1 ----------~----1. Give an instance of the change-making problem for which the greedy algorithm does not yield an optimal solution.
2. Write a pseudocode of the greedy algorithm for the change-making problem
with an amount nand coin denominations d1 > d2 > ... > dm as its input. What
is the time efficiency class of your algorithm?
3. Consider the problem of schedulingn jobs of known durations t1 , t2 , ... , t,. for
execution by a single processor. The jobs can be executed in any order, one job
at a time. You want to find a schedule that minimizes the total time spent by
all the jobs in the system. (The time spent by one job in the system is the sum
of the time spent by this job in waiting plus the time spent on its execution.)
Design a greedy algorithm for this problem. Does the greedy algorithm always
yield an optimal solution?
4. Design a greedy algorithm for the assignment problem (see Section 3.4). Does
your greedy algorithm always yield an optimal solution?
5. Bridge crossing revisited Consider the generalization of the bridge crossing
puzzle (Problem 2 in Exercises l.2) in which we have n > 1 people whose
bridge crossing times are t1 , t2 , ... , t,. All the other conditions of the problem
remain the same: only two people at the time can cross the bridge (and they
move with the speed of the slower of the two) and they must carry with them
the only flashlight the group has.
Design a greedy algorithm for this problem and find how long it will
take to cross the bridge by using this algorithm. Does your algorithm yield a
314
Greedy Technique
minimum crossing time for every instance of the problem? If it does-prove it;
if it does not-find an instance with the smallest number of people for which
this happens.
6. Bachet-Fibonacci weighing problem Find an optimal set ofn weights (wl, w2,
... , wn) so that it would be possible to weigh on a balance scale any integer
load in the largest possible range from 1 to W, provided
a. weights can be put only on the free cup of the scale.
b. weights can be put on both cups of the scale.
7. a. Apply Prim's algorithm to the following graph. Include in the priority
queue all the vertices not already in the tree.
5
b. Apply Prim's algorithm to the following graph. Include in the priority
queue only the fringe vertices (the vertices not in the current tree which
are adjacent to at least one tree vertex).
5
6
6
9
8. The notion of a minimum spanning tree is applicable to a connected weighted
graph. Do we have to check a graph's connectivity before applying Prim's
algorithm or can the algorithm do it by itself?
9. a. How can we use Prim's algorithm to find a spanning tree of a connected
graph with no weights on its edges?
b. Is it a good algorithm for this problem?
10. Prove that any weighted connected graph with distinct weights has exactly
one minimum spanning tree.
11. Outline an efficient algorithm for changing an element's value in a min-heap.
What is the time efficiency of your algorithm?
9.2
Kruskal's Algorithm
315
9.2 Kruskal's Algorithm
In the previous section, we considered the greedy algorithm that "grows" a minimum spanning tree through a greedy inclusion of the nearest vertex to the vertices
already in the tree. Remarkably, there is another greedy algorithm for the minimum spanning tree problem that also always yields an optimal solution. It is
namedKruskal's algorithm [Kru56], after Joseph Kruskal, who discovered the algorithm when he was a second-year graduate student. Kruskal's algorithm looks at
a ntinimum spanning tree for a weighted connected graph G = (V, E) as an acyclic
subgraph with IV I - 1edges for wltich the sum of the edge weights is the smallest.
(It is not difficult to prove that such a subgraph must be a tree.) Consequently,
the algorithm constructs a minimum spanning tree as an expanding sequence of
subgraphs, which are always acyclic but are not necessarily connected on the intermediate stages of the algorithm.
The algorithm begins by sorting the graph's edges in nondecreasing order
of their weights. Then, starting with the empty subgraph, it scans this sorted list
adding the next edge on the list to the current sub graph if such an inclusion does
not create a cycle and simply skipping the edge otherwise.
ALGORITHM
Kruskal(G)
//Kruskal's algorithm for constructing a minimum spanning tree
//Input: A weighted connected graph G = (V, E)
//Output: Er, the set of edges composing a minimum spanning tree of G
sort E in nondecreasing order of the edge weights w(e;) ::S ... ::S w(e,l£ )
1
Er +-- 0; ecounter +-- 0 //initialize the set of tree edges and its size
k +-- 0
//initialize the number of processed edges
while ecounter < lVI- 1 do
k <--k+l
if Er U [e;,} is acyclic
Er +--- ET U {ei);
returnEr
ecounter +--- ecounter + 1
The correctness of Kruskal's algorithm can be proved by repeating the essential steps of the proof of Prim's algorithm given in the previous section. The fact
that Er is actually a tree in Prim's algorithm but generally just an acyclic subgraph
in Kruskal's algorithm turns out to be an obstacle that can be overcome.
Figure 9.4 demonstrates the application of Kruskal's algorithm to the same
graph we used for illustrating Prim's algorithm in Section 9.1. As you trace the
algorithm's operations, note the disconnectedness of some of the intermediate
graphs.
Applying Prim's and Kruskal's algorithms to the same small graph by hand
may create an impression that the latter is simpler than the former. This impression
is wrong because, on each of its iterations, Kruskal's algorithm has to check
whether the addition of the next edge to the edges already selected would create
1
3
Tree edges
6
Illustration
Sorted list of edges
be ef ab bf ef af df ae ed de
1 2 3 4 4 5 5 6 6 8
3
4
4
6
4
4
6
6
be
1
be ef ab bf cf af df ae cd de
1 2 3 4 4 5 5 6 6 8
3
ef
2
be ef ab bf cf af df ae ed de
1234455668
ab
3
be ef ab bf cf af df ae ed de
1234455668
bf
4
be ef ab bf cf af df ae cd de
1234455668
df
5
FIGURE 9.4 Application of Kruskal's algorithm. Selected edges are shown in bold.
316
I
9.2
Kruskal's Algorithm
317
v
(a)
(b)
FIGURE 9.5 New edge connecting two vertices may (a) or may not (b) create a cycle
a cycle. It is not difficult to see that a new cycle is created if and only if the new
edge connects two vertices already connected by a path, i.e., if and only if the two
vertices belong to the same connected component (Figure 9.5). Note also that
each connected component of a subgraph generated by Kruskal's algorithm is a
tree because it has no cycles.
In view of these observations, it is convenient to use a slightly different
interpretation of Kruskal's algorithm. We can consider the algorithm's operations
as a progression through a series of forests containing all the vertices of a given
graph and some of its edges. The initial forest consists of IV 1 trivial trees, each
comprising a single vertex of the graph. The final forest consists of a single tree,
which is a minimum spanning tree of the graph. On each iteration, the algorithm
takes the next edge (u, v) from the sorted list of the graph's edges, finds the trees
containing the vertices u and v, and, if these trees are not the same, unites them
in a larger tree by adding the edge (u, v).
Fortunately, there are efficient algorithms £or doing so, including the crucial
check whether two vertices belong to the same tree. They are called union-find
algorithms. We discuss them in the following subsection. With an efficient unionfind algorithm, the running time of Kruskal's algorithm will be dominated by
the time needed for sorting the edge weights of a given graph. Hence, with an
efficient sorting algorithm, the time efficiency of Kruskal's algorithm will be in
O(IEilog lEI).
Disjoint Subsets and Union-Find Algorithms
Kruskal's algorithm is one of a number of applications that require a dynamic
partition of some n-element setS into a collection of disjoint subsets Sl> S2, ... , Sk.
After being initialized as a collection of n one-element subsets, each containing
a different element of S, the collection is subjected to a sequence of intermixed
union and find operations. (Note that the number of union operations in any such
sequence must be bounded above by n - 1 because each union increases a subset's
size at least by 1 and there are only n elements in the entire setS.) Thus, we are
318
Greedy Technique
dealing here with an abstract data type of a collection of disjoint subsets of a finite
set with the following operations:
makeset(x) creates a one-element set {x ].It is assumed that this operation can
be applied to each of the elements of set S only once;
find(x) returns a subset containing x;
union(x, y) constructs the union of the disjoint subsets Sx and Sy containing
x andy, respectively, and adds it to the collection to replace Sx and Sy, which
are deleted from it.
For example, letS= {1, 2, 3, 4, 5, 6]. Then makeset(i) creates the set {i] and
applying this operation six times initializes the structure to the collection of six
singleton sets:
{1], {2], {3}, {4], {5], {6].
Performing union(l, 4) and union(5, 2) yields
{1, 4], {5, 2], {3], {6],
and, if followed by union(4, 5) and then by union(3, 6), we end up with the disjoint
subsets
{1, 4, 5, 2], {3, 6].
Most implementations of this abstract data type use one element from each of
the disjoint subsets in a collection as that subset's representative. Some implementations do not impose any specific constraints on such a representative; others do
so by requiring, say, the smallest element of each subset to be used as the subset's
representative. Also, it is usually assumed that set elements are (or can be mapped
into) integers.
There are two principal alternatives for implementing this data structure. The
first one, called the quick find, optimizes the time efficiency of the fmd operation;
the second one, called the quick union, optimizes the union operation.
The quick find uses an array indexed by the elements of the underlying set
S; the array's values indicate the representatives of the subsets containing those
elements. Each subset is implemented as a linked list whose header contains the
pointers to the first and last elements of the list along with the number of elements
in the list (see Figure 9.6 for an example).
Under this scheme, the implementation of makeset(x) requires assigning the
corresponding element in the representative array to x and initializing the corre-
sponding linked list to a single node with the x value. The time efficiency of this
operation is obviously in 8 (1), and hence the initialization of n singleton subsets is
in EJ(n). The efficiency ofjind(x) is also in 8(1): all we need to do is to retrieve the
x's representative in the representative array. Executing union(x, y) takes longer.
A straightforward solution would simply append the y's list to the end of the x's
list, update the information about their representative for all the elements in the
9.2 Kruskal's Algorithm
319
size last first
list 1
I 4 1r 1·1 .I
list
21 0 Inull! null[
list
31 2
list
41 0 [null! null!
list sJ 0
list
61 0
1·1
·I 4 1·1 ·I
I r 1·1 ·I 3 1·1 ·I 6f Inull!
Inull! null!
Inull! null!
5
·I t2 Inull[
1·1
subset representatives
element index
representative
1
2
3
4
5
6
1
1
3
1
1
3
FIGURE 9.6 Linked-list representation of subsets (1. 4, 5, 2) and (3, 6) obtained by quick
find after performing union(1, 4), union(5, 2), union(4, 5), and union(3, 6).
The lists of size 0 are considered deleted from the collection.
y list, and then delete the y's list from the collection. It is easy to verify, however,
that with this algorithm the sequence o£ union operations
union(2, 1), union(3, 2), ... , union(i + 1, i), ... , union(n, n- 1)
runs in E>(n 2) time, which is slow compared with several known alternatives.
A simple way to improve the overall efficiency of a sequence of union operations is to always append the shorter of the two lists to the longer one, with
ties broken arbitrarily. Of course, the size of each list is assumed to be available
by, say, storing the number of elements in the list's header. This modification is
called the union by size. Though it does not improve the worst-case efficiency of
a single application of the union operation (it is still in E>(n) ), the worst-case running time of any legitimate sequence of union-by-size operations turns out to be
in O(n log n) 2
Here is a proof of this assertion. Let a; be an element of set S whose disjoint
subsets we manipulate and let Ai be the number of times ai's representative is
2.
This is a specific example of the usefulness of the amortized efficiency we mentioned back in Chapter 2.
The time efficiency of any sequence of n union-by-size operations is more efficient than the worst-case
efficiency of its single application times n.
--,
'
320
Greedy Technique
4
4
5
2
(a)
(b)
FIGURE 9.7 (a) Forest representation of subsets [1, 4, 5, 2) and [3, 6) used by quick
union. (b) Result of union(S, 6).
updated iu a sequence of union-by-size operations. How large can A; get if setS
has n elements? Each time ai 's representative is updated, a 1 must be in a smaller
subset involved in computing the union whose size will be at least twice as large as
the size of the subset containing a,. Hence, when a, 's representative is updated for
the first time, the resulting set will have atleast two elements; when it is updated for
the second time, the resulting set will have at least four elements; and, in general,
if it is updated A, times, the resulting set will have at least zA; elements. Since the
entire set S has n elements, 2A; .:::=: n and hence Ai ::: log2 n. Therefore, the total
number of possible updates of the representatives for all n elements in S will not
exceed n log2 n.
Thus, for union by size, the time efficiency of a sequence of at most n - 1
unions and m finds is in O(n log n + m).
The quick union~the second principal alternative for implementing disjoint
subsets~represents each subset by a rooted tree. The nodes of the tree contain
the subset's elements (one per node), with the root's element considered the
subset's representative; the tree's edges are directed from children to their parents
(Figure 9.7). (In addition, a mapping of the set elements to their tree nodes~
implemented, say, as an array of pointers~is maintained. This mapping is not
shown in Figure 9.7 for the sake of simplicity.)
For this implementation, makeset(x) requires the creation of a single-node
tree, which is a E>(l) operation; hence the initialization of n singleton subsets is in
E>(n). A union(x, y) is implemented by attaching the root of the y's tree to the root
of the x 's tree (and deleting the y's tree from the collection by making the pointer
to its root null). The time efficiency of this operation is clearly 8(1). Afind(x) is
performed by following the pointer chain from the node containing x to the tree's
root (whose element is returned as the subset's representative). Accordingly, the
time efficiency of a single find operation is in 0 (n) because a tree representing a
subset can degenerate into a linked list with n nodes.
This time bound can be improved. The straightforward way for doing so is to
always perform a union operation by attaching a smaller tree to the root of a larger
9.2
Kruskal's Algorithm
321
FIGURE 9.8 Path compression
one, with ties broken arbitrarily. The size of a tree can be measured either by the
number of nodes (this version is called union by size) or by its height (this version
is called union by rank). Of course, these options require storing, for each node
of the tree, either the number of node descendants or the height of the subtree
rooted at that node, respectively. One can easily prove that in either case the
height of the tree will be logarithmic, making it possible to execute each find in
0 (log n) time. Thus, for quick union, the time efficiency of a sequence of at most
n- 1 unions and m finds is in O(n + m log n).
In fact, an even better efficiency can be obtained by combining either variety of quick union with path compression. This modification makes every node
encountered during the execution of a fmd operation point to the tree's root (Figure 9.8).
According to a quite sophisticated analysis that goes beyond the level of
this book (see (Tar84]), this and similar techniques improve the efficiency of a
sequence of at most n - 1 unions and m finds to only slightly worse than linear.
-----Exercises 9.2 - - - - - - - - - - - - - - - 1. Apply Kruskal's algorithm to find a minimum spanning tree of the following
graphs.
a.
~
6
2
+
322
Greedy Technique
b.
2. Indicate whether the following statements are true or false:
a. If e is a minimum-weight edge in a connected weighted graph, it must be
among edges of at least one minimum spanning tree of the graph.
· b. If e is a minimum-weight edge in a connected weighted graph, it must be
among edges of each minimum spanning tree of the graph.
c. If edge weights of a connected weighted graph are all distinct, the graph
must have exactly one minimum spanning tree.
d. If edge weights of a connected weighted graph are not all distinct, the graph
must have more than one minimum spanning tree.
/;
(_3yWhat changes, if any, need to be made in algorithm Kruskal to make it find
a minimum spanning forest for an arbitrary graph? (A minimum spanning
forest is a forest whose trees are minimum spanning trees of the graph's
connected components.)
4. Will either Kruskal's or Prim's algorithm work correctly on graphs that have
negative edge weights?
5. Design an algorithm for finding a maximum spanning tree-a spanning tree
with the largest possible edge weight-of a weighted connected graph.
6. Rewrite the pseudocode of Kruskal's algorithm in terms of the operations of
the disjoint subsets' ADT.
7. Prove the correctness of Kruskal's algorithm.
8. Prove that the time efficiency offind(x) is in O(logn) for the union-by-size
version of quick union.
9. Find at least two Web sites with animations ofKruskal's and Prim's algorithms.
Discuss their merits and demerits.
10. Design and conduct an experiment to empirically compare the efficiencies
-""' of Prim's and Kruskal's algorithms on random graphs of different sizes and
densities.
9.3 Dijkstra's Algorithm
323
11. Steiner tree Four villages are located at the vertices o£ a unit square in the
Euclidean plane. You are asked to connect them by the shortest network of
roads so that there is a path between every pair of the villages along those
roads. Find such a network.
9.3 Dijkstra's Algorithm
In this section, we consider the single-source shortest-paths problem: for a given
vertex called the source in a weighted connected graph, find shortest paths to all
its other vertices. It is important to stress that we are not interested here in a
single shortest path that starts at the source and visits all the other vertices. This
would have been a much more difficult problem (actually, a version of the traveling
salesman problem mentioned in Section 3.4 and discussed again later in the book).
The single-source shortest-paths problem asks for a family of paths, each leading
from the source to a different vertex in the graph, though some paths may, of
course, have edges in common.
A variety of practical applications of the shortest-paths problem have made
the problem a very popular object o£ study. There are several well-known algorithms for solving it, including Floyd's algorithm for the more general all-pairs
shortest-paths problem discussed in Chapter 8. Here, we consider the best-known
algorithm for the single-source shortest-paths problem, called Dijkstra's algorithm3 This algorithm is applicable to graphs with nonnegative weights only. Since
in most applications this condition is satisfied, the limitation has not impaired the
popularity of Dijkstra's algorithm.
Dijkstra's algorithm finds the shortest paths to a graph's vertices in order of
their distance from a given source. First, it finds the shortest path from the source
to a vertex nearest to it, then to a second nearest, and so on. In general, before its
ith iteration commences, the algorithm has already identified the shortest paths
to i - 1 other vertices nearest to the source. These vertices, the source, and the
edges of the shortest paths leading to them from the source form a subtree T,
of the given graph (Figure 9.9). Since all the edge weights are nonnegative, the
next vertex nearest to the source can be found among the vertices adjacent to the
vertices of 1j. The set of vertices adjacent to the vertices in Ti can be referred to as
"fringe vertices"; they are the candidates from which Dijkstra's algorithm selects
the next vertex nearest to the source. (Actually, all the other vertices can be treated
as fringe vertices connected to tree vertices by edges of infinitely large weights.)
To identify the ith nearest vertex, the algorithm computes, for every fringe vertex
3.
Edsger W. Dijkstra (1930-2002), a noted Dutch pioneer of the science and industry of computing,
discovered this algroithm in the mid-1950s. Dijkstra said of his algorithm: "This was the first graph
problem I ever posed myself and solved. The amazing thing was that I didn't publish it. It was not
amazing at the time. At the time, algorithms were hardly considered a scientific topic.''
324
Greedy Technique
FIGURE 9.9 Idea of Dijkstra's algorithm. The subtree of the shortest paths already found
is shown in bold. The next nearest to the source v0 vertex, u*, is selected
by comparing the lengths of the subtree's paths increased by the distances
to vertices adjacent to the subtree's vertices.
u, the sum of the distance to the nearest tree vertex v (given by the weight of the
edge (v, u)) and the length d., of the shortest path from the source to v (previously
determined by the algorithm) and then selects the vertex with the smallest such
sum. The fact that it suffices to compare the lengths of such special paths is the
central insight of Dijkstra's algorithm.
To facilitate the algorithm's operations, we label each vertex with two labels.
The numeric label d indicates the length of the shortest path from the source to
this vertex found by the algorithm so far; when a vertex is added to the tree, d
indicates the length of the shortest path from the source to that vertex. The other
label indicates the name of the next-to-last vertex on such a path, i.e., the parent of
the vertex in the tree being constructed. (It can be left unspecified for the source
s and vertices that are adjacent to none of the current tree vertices.) With such
labeling, finding the next nearest vertex u* becomes a simple task of finding a
fringe vertex with the smallest d value. Ties can be broken arbitrarily.
After we have identified a vertex u* to be added to the tree, we need to perform
two operations:
"'
"'
Move u* from the fringe to the set of tree vertices.
For each remaining fringe vertex u that is connected to u* by an edge of
weight w(u*, u) such that d.,.+ w(u*, u) <d.,, update the labels of u by u*
and du* + w(u*, u), respectively.
Figure 9.10 demonstrates the application of Dijkstra's algorithm to a specific
graph.
The labeling and mechanics of Dijkstra 's algorithm are quite similar to those
used by Prim's algorithm (see Section 9.1). Both of them construct an expanding
subtree of vertices by selecting the next vertex from the priority queue of the
remaining vertices. It is important not to mix them up, however. They solve
9.3
Tree vertices
Dijkstra's Algorithm
Remaining vertices
a(-, 0)
b(a, 3) c(-, oo) d(a, 7) e(-, oo)
b(a, 3)
c(b, 3+4) d(b, 3+2) e(-, oo)
d(b, 5)
c(b, 7) e(d, 5 + 4)
c(b, 7)
e(d, 9)
325
Illustration
e(d, 9)
The shortest paths (identified by following nonnumeric labels
backward from a destination vertex in the left column to the
source) and their lengths (given by numeric labels of the tree
vertices) are
fromatob:
fromatod:
from a to c:
from a toe:
a-b
a-b-d
a-b-c
a-b-d-e
of length 3
of length 5
of length 7
of length 9
FIGURE 9.10 Application of Dijkstra's algorithm. The next closest vertex is shown in
bold.
326
Greedy Technique
different problems and therefore operate with priorities computed in a different
manner: Dijkstra's algorithm compares path lengths and therefore must add edge
weights, while Prim's algorithm compares the edge weights as given.
Now we can give a pseudocode of Dijkstra's algorithm. It is spelled outin more detail than Prim's algorithm was in Section 9.1-in terms of explicit
operations on two sets oflabeled vertices: the set VT of vertices for which a shortest
path has already been found and the priority queue Q of the fringe vertices. (Note
that in the following pseudocode, Vr contains a given source vertex and the fringe
contains the vertices adjacent to it after iteration 0 is completed.)
ALGORITHM
Dijkstra(G, s)
//Dijkstra's algorithm for single-source shortest paths
//Input: A weighted connected graph G = (V, E) with nonnegative weights
II
and its vertex s
//Output: The length d" of a shortest path from s to v
II
and its penultimate vertex p" for every vertex v in V
lnitialize(Q) //initialize vertex priority queue to empty
for every vertex v in V do
dv +- oo;
Pv +-- null
Insert(Q, v, d") //initialize vertex priority in the priority queue
d, <-- 0; Decrease(Q, s, d,) //update priority of s with d,
Vr <-- 0
fori <-Oto IVI-1do
u* <-- DeleteMin(Q) //delete the minimum priority element
Vr <-- Vr U {u*)
for every vertex u in V - Vr that is adjacent to u* do
ifdu*+w(u*, u) <du
du +-- du* + w(u*, u);
Decrease(Q, u, d,)
p 11 +-- u*
The time efficiency of Dijkstra's algorithm depends on the data structures
used for implementing the priority queue and for representing an input graph itself. For the reasons explained in the analysis of Prim's algorithm in Section 9.1, it
is in 8(1 V 12) for graphs represented by their weight matrix and the priority queue
implemented as an unordered array. For graphs represented by their adjacency
lists and the priority queue implemented as a min-heap, it is in O(IEilog lVI). A
still better upper bound can be achieved for both Prim's and Dijkstra's algorithms
if the priority queue is implemented using a sophisticated data structure called
the Fibonacci heap (e.g., (Wei98]). However, its complexity and a considerable
overhead make such an improvement primarily of theoretical value.
9.3
Dijkstra's Algorithm
327
-----Exercises 9"3 - - - - - - - - - - - - - - - 1. Explain what adjustments, if any, need to be made in Dijkstra's algorithm
and/or in an underlying graph to solve the following problems.
a. Solve the single-source shortest-paths problem for directed weighted
graphs.
b. Find a shortest path between two given vertices of a weighted graph or
digraph. (This variation is called the single-pair shortest-path problem.)
c. Find the shortest paths to a given vertex from each other vertex of a
weighted graph or digraph. (This variation is called the single-destination
shortest-paths problem.)
d. Solve the single-source shortest-paths problem in a graph with nonnegative
numbers assigned to its vertices (and the length of a path defined as the sum
of the vertex numbers on the path).
2. Solve the following instances of the single-source shortest-paths problem with
vertex a as the source:
a.
b.
c-1~
5
6
2
c
d
6
h
f
4
5
3
2
e
6
5
3
9
3. Give a counterexample that shows that Dijkstra's algorithm may not work for
a weighted connected graph with negative weights.
4. Let T be a tree constructed by Dijkstra"s algorithm in the process of solving
the single-source shortest-paths problem for a weighted connected graph G.
a. True or false: Tis a spanning tree of G?
b. True or false: T is a minimum spanning tree of G?
I
328
Greedy Technique
5. Write a pseudocode of a simpler version of Dijkstra's algorithm that finds
only the distances (i.e., the lengths of shortest paths but not shortest paths
themselves) from a given vertex to all other vertices of a graph represented
by its weight matrix.
6. Prove the correctness of Dijkstra's algorithm for graphs with positive weights.
Q
Design a linear-time algorithm for solving the single-source shortest -paths
problem for dags (directed acyclic graphs) represented by their adjacency lists.
Design an efficient algorithm for finding the length of a longest path in a dag.
(This problem is important because it determines a lower bound on the total
time needed for completing a project composed of precedence-constrained
tasks.)
9. Shortest-path modeling Assume that you have a model of a weighted connected graph made of balls (representing the vertices) connected by strings
of appropriate lengths (representing the edges).
a. Describe how you can solve the single-pair shortest -path problem with this
model.
b. Describe how you can solve the single-source shortest-paths problem with
this model.
10. Revisit Problem 6 in Exercises 1.3 about determining the best route for a
subway passenger to take from one designated station to another in a welldeveloped subway system like those in Washington, DC, and London, UK.
Write a program for this task.
9.4 Huffman Trees
Suppose we have to encode a text that comprises characters from some n-character
alphabet by assigning to each of the text's characters some sequence of bits called
the codeword. For example, we can use a fixed-length encoding that assigns to
each character a bit string of the same length m (m 2: log2 n ). This is exactly what
the standard ASCII code does. One way of getting a coding scheme that yields
a shorter bit string on the average is based on the old idea of assigning shorter
codewords to more frequent characters and longer codewords to less frequent
characters. (This idea was used, in particular, in the telegraph code invented in
the mid-19th century by Samuel Morse. In that code, frequent letters such as e
O and a (--) are assigned short sequences of dots and dashes while infrequent
letters such as q (- - ·-) and z (- - .. ) have longer ones.)
Variable-length encoding, which assigns codewords of different lengths
to different characters, introduces a problem that fixed-length encoding does
not have. Namely, how can we tell how many bits of an encoded text represent the first (or, more generally, the ith) character? To avoid this complication, we can limit ourselves to prefix-free (or simply prefix) codes. In a prefix
l
9.4
Huffman Trees
329
code, no codeword is a prefix of a codeword of another character. Hence, with
such an encoding, we can simply scan a bit string until we get the first group
of bits that is a codeword for some character, replace these bits by this character,
and repeat this operation until the bit string's end is reached.
If we want to create a binary prefix code for some alphabet, it is natural to
associate the alphabet's characters with leaves of a binary tree in which all the
left edges are labeled by 0 and all the right edges are labeled by 1 (or vice versa).
The codeword of a character can then be obtained by recording the labels on the
simple path from the root to the character's leaf. Since there is no simple path
to a leaf that continues to another leaf, no codeword can be a prefix of another
codeword; hence, any such tree yields a prefix code.
Among the many trees that can be constructed in this manner for a given
alphabet with known frequencies of the character occurrences, how can we construct a tree that would assign shorter bit strings to high-frequency characters and
longer ones to low-frequency characters? It can be done by the following greedy
algorithm, invented by David Huffman while he was a graduate student at MIT
[Hu£52].
Huffman's Algorithm
Step 1 Initialize n one-node trees and label them with the characters of the
alphabet. Record the frequency of each character in its tree's root to
indicate the tree's weight. (More generally, the weight of a tree will be
equal to the sum of the frequencies in the tree's leaves.)
Step 2 Repeat the following operation until a single tree is obtained. Find two
trees with the smallest weight (ties can be broken arbitrarily, but see
Problem 2 in the exercises). Make them the left and right subtree of
a new tree and record the snm of their weights in the root of the new
tree as its weight.
A tree constructed by the above algorithm is called a Huffman tree. It
defines-in the manner described-a Huffman code.
EXAMPLE Consider the five-character alphabet {A, B, C, D, _}with the following
occurrence probabilities:
character
A
B
c
D
probability
0.35
0.1
0.2
0.2
0.15
The Huffman tree construction for this input is shown in Figure 9.11. The
resulting codewords are as follows:
character
A
B
c
D
probability
codeword
0.35
11
0.1
100
0.2
00
0.2
01
0.15
101
330
Greedy Technique
·s
LJ
3
! '
'~.
Eij 10~51 I 0~21 I o: I 10:51
I 0~21 I 0~21
10:51
0.6
0
,~p
FIGURE 9.11 Example of constructing a Huf1man codiRg tree
Hence, DAD is encoded as 011101, and 10011011011101 is decoded as BAD_AD.
With the occurrence probabilities given and the codeword lengths obtained,
the expected number of bits per character in this code is
I
I
2. 0.35 + 3. 0.1 + 2. 0.2 + 2. 0.2 + 3. 0.15 = 2.25.
I
9.4 Huffman Trees
331
Had we used a fixed-length encoding for the same alphabet, we would have to
use at least three bits per each character. Thus, for this toy example, Huffman's
code achieves the compression ratio-a standard measure of a compression algorithm's effectiveness-of (3- 2.25)/3 · 100% = 25%. In other words, we should
expect that Huffman's encoding of a text will use 25% less memory than its fixedlength encoding. (Extensive experiments with Huffman codes have shown that the
compression ratio for tbis scheme typically falls between 20% and 80%, depending
Ill
on the characteristics of the text being compressed.)
Huffman's encoding is one of the most important file compression methods. In
addition to its simplicity and versatility, it yields an optimal, i.e., minimal-length,
encoding (provided the probabilities of character occurrences are independent
and known in advance). The simplest version of Huffman compression calls, in
fact, for a preliminary scanning of a given text to count the frequencies of character
occurrences in it. Then these frequencies are used to construct a Huffman coding
tree and encode the text as described above. This scheme makes it necessary,
however, to include the information about the coding tree into the encoded text
to make its decoding possible. This drawback can be overcome by using dynamic
Huffman encoding, in which the coding tree is updated each time a new character
is read from the source text (see, e.g., [SayOO]).
It is important to note that applications of Huffman's algorithm are not limited
to data compression. Suppose we have n positive numbers wb w2 , ... , w that
11
have to be assigned to 11 leaves of a binary tree, one per node. If we define the
weighted path length as the sum I:;'~ 1 liw 1 , where li is tbe length of the simple
path from the root to the ith leaf, how can we construct a binary tree with
minimum weighted path length? It is this more general problem that Huffman's
algorithm actually solves. (For the coding application, li and w 1 are the length of
the codeword and the frequency of the ith character, respectively.) This problem
arises in many situations involving decision making. Consider, for example, the
game of guessing a chosen object from 11 possibilities (say, an integer between 1
and 11) by asking questions answerable by yes or no. Different strategies for playing
this game can be modeled by decision trees 4 such as those depicted in Figure 9.12
forn =4.
The length of the simple path from the root to a leaf in such a tree is equal
to the number of questions needed to get to the chosen number represented by
the leaf. If number i is chosen with probability Pi· the sum I:;'~ liPi• where li is
the length of the simple path from the root to the ith leaf, indicates the average
number of questions needed to "guess" the chosen number witb a game strategy
represented by its decision tree. If each of the numbers is chosen with the same
probability of 1/n, the best strategy is to successively eliminate half (or almost half)
the candidates as binary search does. This may not be the case for arbitrary Pi's,
1
4.
Decision trees are discussed in more detail in Section 11.2.
332
Greedy Technique
n=4
no
n=3
n>3
no
yes
yes
no
yes
FIGURE 9.12 Two decision trees for guessing an integer between 1 and 4
however. (For example. if n = 4 and Pt = 0.1. P2 = 0.2. p 3 = 0.3. and P4 = 0.4. the
minimum weighted path tree is the rightmost one in Figure 9.12.) Thus. we need
Huffman's algorithm to solve this problem in its general case.
In conclusion, it is worthwhile to remember that this is the second time we
are encountering the problem of constructing an optimal binary tree. In Section
8.3, we discussed the problem of constructing an optimal binary search tree with
positive numbers (the search probabilities) assigned to every node of the tree. In
this section, given numbers are assigned just to leaves. The latter problem turns
out to be easier: it can be solved by the greedy algorithm whereas the former is
solved by the more complicated dynamic programming algorithm.
-----Exercises 9 . 4 - - - - - - - - - - - - - - - 1. a. Construct a Huffman code for the following data:
character
A
B
c
D
probability
0.4
0.1
0.2
0.15
0.15
b. Encode the text ABACABAD using the code of question (a).
c. Decode the text whose encoding is 100010111001010 in the code of question (a).
2. For data transmission purposes, it is often desirable to have a code with a
minimum variance of the codeword lengths (among codes of the same average
length). Compute the average and variance of the codeword length in two
Huffman codes that result from a different tie breaking during a Huffman
code construction for the following data:
character
A
B
c
D
E
probability
0.1
0.1
0.2
0.2
0.4
Summary
333
3. Indicate whether each of the following properties are true for every Huffman
code.
a. The codewords of the two least frequent characters have the same length.
b. The codeword's length of a more frequent character is always smaller than
or equal to the codeword's length of a less frequent one.
4. What is the maximal length of a codeword possible in a Huffman encoding of
an alphabet of n characters?
5. a. Write a pseudocode for the Huffman-tree construction algorithm.
b. What is the time efficiency class of the algorithm for constructing a Huffman tree as a function of the alphabet's size?
! 6. Show that a Huffman tree can be constructed in li11ear ti111e if the alphabet's
--- 'characters are given in a sorted order of their freqUe~n·ci;;s:- ~-
'•, 7J Given a Huffman coding tree, wbich algorithm would you use to get the
codewords for all the characters? What is its time-efficiency class as a function
of the alphabet's size?
s:' ofExplain
how one can generate a Huffman code without an explicit generation
a Huffman coding tree.
9. a. Write a program that constructs a Huffman code for a given English text
and encode it.
b. Write a program for decoding an English text that has been encoded with
a Huffman code.
c. Experiment with your encoding program to find a range of typical compression ratios for Huffman's encoding of English texts of, say, 1000 words.
d. Experiment with your encoding program to find out how sensitive the
compression ratios are to using standard estimates o£ frequencies instead
of actual frequencies of character occurrences in English texts.
10. Card guessing Design a strategy that minimizes the expected number of
questions asked in the following game [Gar94], #52. You have a deck of cards
that consists of one ace of spades, two deuces of spades, three threes, and on up
to nine nines, making 45 cards in all. Someone draws a card from the shuffled
deck, which you have to identify by asking questions answerable with yes or
no.
SUMMARY
I!
The greedy technique suggests constructing a solution to an optimization
problem through a sequence of steps, each expanding a partially constructed
solution obtained so far, until a complete solution to the problem is reached.
334
Greedy Technique
On each step, the choice made must be feasible, locally optimal, and
irrevocable.
"
Prim's algorithm is a greedy algorithm for constructing a minimum spanning
tree of a weighted connected graph. It works by attaching to a previously
constructed subtree a vertex closest to the vertices already in the tree.
101
Kruskal's algorithm is another greedy algorithm for the minimum spanning
tree problem. It constructs a minimum spanning tree by selecting edges in
increasing order of their weights provided that the inclusion does not create a
cycle. Checking the latter condition efficiently requires an application of one
of the so-called union-find algorithms.
101
Dijkstra's algorithm solves the single-source shortest-paths problem of finding
shortest paths from a given vertex (the source) to all the other vertices of a
weighted graph or digraph. It works as Prim's algorithm but compares path
lengths rather than edge lengths. Dijkstra's algorithm always yields a correct
solution for a graph with nonnegative weights.
101
A Huffman tree is a binary tree that minimizes the weighted path length from
the root to the leaves of predefined weights. The most important application
of Huffman trees are Huffman codes.
101
A Huffman code is an optimal prefix-free variable-length encoding scheme
that assigns bit strings to characters based on their frequencies in a given
text. This is accomplished by a greedy construction of a binary tree whose
leaves represent the alphabet characters and whose edges are labeled with
O's and l's.
1
Iterative Improvement
The most successful men in the end are those whose success is the result
of steady accretion.
-Alexander Graham Bell (1835-191 01
T
he greedy strategy, considered in the preceding chapter, constructs a solution
to an optimization problem piece by piece, always adding a locally optimal
piece to a partially constructed solution. In this section, we discuss a different
approach to designing algorithms for optimization problems. It starts with some
feasible solution (a solution that satisfies all the constraints of the problem) and
proceeds to improve it by repeated applications of some simple step. This step
typically involves a small, localized change yielding a feasible solution with an
improved value of the objective function. When no such change improves the
value of the objective function, the algorithm returns the last feasible solution as
optimal and stops.
There can be several obstacles to the successful implementation of this idea.
First, we need an initial feasible solution. For some problems, we can always start
with a trivial solution or use an approximate solution obtained by some other (e.g.,
greedy) algorithm. But for others, finding an initial solution may require as much
effort as solving the problem after a feasible solution has been identified. Second,
it is not always clear what changes should be allowed in a feasible solution so that
we can check efficiently whether the current solution is locally optimal and, if not,
replace it with a better one. Third-and this is the most fundamental difficultyis an issue of local versus global extremum (maximum or minimum). Think about
the problem of finding a highest point in a hilly area with no map on a foggy day.
A logical thing to do would be to start walking "up the hill" from the point you are
at until it becomes impossible to do so because no direction would lead up. You
will have reached a local highest point, but because of a limited feasibility, there
will be no simple way to tell whether the point is the highest (global maximum
you are after) in the entire area.
Fortunately, there are important problems that can be solved by iterativeimprovement algorithms. The most important of them is linear programming.
335
336
Iterative Improvement
We have already encountered this topic in Section 6.6. Here, in Section 10.1,
we introduce the simplex method, the classic algorithm for linear programming.
Discovered by the U.S. mathematician George B. Dantzig in 1947, this algorithm
has proved to be one of the most consequential achievements in the history of
algorithmics.
In Section 10.2, we consider the important problem of maximizing the amount
of flow that can be sent through a network with links of limited capacities. Tbis
problem is a special case of linear programming. However, its special structure
makes it possible to solve the problem by algorithms that are more efficient than
the simplex method. We outline the classic iterative improvement algorithm for
this problem, discovered by the American mathematicians L. R. Ford, Jr., and D.
R. Fulkerson in the 1950s.
The last two sections of the chapter deal with bipartite matching. This is
a problem of finding an optimal pairing of elements taken from two disjoint
sets. Examples include matching workers and jobs, high school graduates and
colleges, and men and women for marriage. Section 10.3 deals with the problem
of maximizing the number of matched pairs; Section 10.4 is concerned with the
matching stability.
We also discuss several iterative improvement algorithms in Section 12.3,
where we consider approximation algorithms for the traveling salesman and knapsack problems. Other examples of iterative improvement algorithms can be found
in the algorithms textbook by Moret and Shapiro [Mor91], books on continuous
and discrete optimization (e.g., [Nem89]), and the literature on heuristic search
(e.g., [MicOO]).
10.1 The Simplex Method
We have already encountered linear programming (see Section 6.6)-the general
problem of optimizing a linear function of several variables subject to a set of
linear constraints:
+ ... + C X
maximize (or minimize)
c1x1
subject to
ai1x 1
11
11
+ ... + llj X ::S (or > or =) bi fori= 1, ... , m
11
11
Xt 2:. 0, ... , X 11 ::: 0.
(16.1)
We mentioned there that many important practical problems can be modeled as
instances of linear programming. Two researchers, L. V. Kantorovich of the former
Soviet Union and the Dutch-American T. C. Koopmans, were even awarded the
Nobel Prize in 1975 for their contributions to linear programming theory and
its applications to economics. Apparently because there is no Nobel Prize in
mathematics, the Royal Swedish Academy of Sciences failed to honor the U.S.
mathematician G. B. Dantzig, who is universally recognized as the father of linear
10.1
The Simplex Method
337
programming in its modern form and the inventor of the simplex method, the
classic algorithm for solving such problems. 1
Geometric Interpretation of linear Programming
Before we introduce a general method for solving linear programming problems,
let us consider a small example, which will help us to see the fundamental properties of such problems.
EXAMPLE 1 Consider the following linear programming problem in two variables:
maximize 3x + 5y
subject to x + y :" 4
X+ 3y :<( 6
X ": 0,
(10.2)
y ": 0.
By definition, a feasible solution to this problem is any point (x, y) that satisfies
all the constraints of the problem; the problem's feasible region is the set of all
its feasible points. It is instructive to sketch the feasible region in the Cartesian
plane. Recall that any equation ax + by = c, where coefficients a and b are not
both equal to zero, defines a straight line. Such a line divides the plane into two
half-planes: for all the points in one of them, ax+ by < c, while for all the points
in the other, ax +by > c. (It is easy to determine which of the two half-planes
is which: take any point (x0 , y0 ) not on the line ax + by = c and check which of
the two inequalities hold, ax 0 + by0 >cor ax0 + by0 <c.) In particular, the set of
points defined by inequality x + y :" 4 comprises the points on and below the line
x + y = 4, and the set of points defined by inequality x + 3y :" 6 comprises the
points on and below the line x + 3y = 6. Since the points of the feasible region
must satisfy all the constraints of the problem, the feasible region is obtained by
the intersection of these two half-planes and the first quadrant of the Cartesian
plane defined by the nonnegativity constraints x ::> 0, y ::> 0 (see Figure 10.1). Thus,
the feasible region for problem (10.2) is the convex polygon with the vertices (0,0),
(4,0), (0,2), and (3,1). (The last point, which is the point of intersection of the lines
x + y = 4 and x + 3y = 6, is obtained by solving the system of these two linear
equations.) Our task is to find an optimal solution, a point in the feasible region
with the largest value of the objective function z = 3x + Sy.
Are there feasible solutions for which the value of the objective function
equals, say, 20? The points (x, y) for which the objective function z = 3x + Sy is
1.
George B. Dantzig (1914--2005) has received many honors, including the National Medal of Science
presented by the president of the United States in 1976. The citation states that the National Medal was
awarded "for inventing linear programming and discovering methods that led to wide-scale scientific
and technical applications to important problems in logistics, scheduling, and network optimization,
and to the use of computers in making efficient use of the mathematical theory."
-.....,
I!
338
Iterative Improvement
y
X+ 3y= 6
X
X+ y= 4
FIGURE 10.1 Feasible region of problem (1 0.2)
equal to 20 form the line 3x + Sy = 20. Since this line does not have common points
with the feasible region-see Figure 10.2-the answer to the posed question is no.
On the other hand, there are infinitely many feasible points for which the objective
function is equal to, say, 10: they are the intersection points of the line 3x + Sy = 10
with the feasible region. Note that the lines 3x + Sy = 20 and 3x + Sy = 10 have
the same slope, as would any line defined by equation 3x + Sy = z where z is
some constant. Such lines are called level lines of the objective function. Thus,
our problem can be restated as finding the largest value of the parameter z for
which the level line 3x + Sy = z has a common point with the feasible region.
We can find this line either by shifting, say, the line 3x + Sy = 20 south-west
(without changing its slope!) toward the feasible region until it hits the region for
the first time or by shifting, say, the line 3x + Sy = 10 north-east until it hits the
feasible region for the last time. Either way, it will happen at the point (3, 1) with
the corresponding z value 3 · 3 + 5 · 1 = 14. This means that the optimal solution
to the linear programming problem in question is x = 3, y = 1, with the maximal
value of the objective function equal to 14.
Note that if we had to maximize z = 3x + 3y as the objective function in problem (10.2), the level line 3x + 3y = z for the largest value of z would coincide with
the boundary line segment that has the same slope as the level lines (draw this
line in Figure 10.2). Consequently, all the points of the line segment between vertices (3,1) and ( 4,0), including the vertices themselves, would be optimal solutions,
yielding, of course, the same maximal value of the objective function.
II
10.1
The Simplex Method
339
v
X
3x+5y=14
3x+5y=10
FIGURE 10.2 Solving a two-dimensional linear programming problem geometrically
Does every linear programming problem have an optimal solution that can
be found at a vertex of its feasible region? Without appropriate qualifications,
the answer to this question is no. To begin with, the feasible region of a linear
programming problem can be empty. For example, if the constraints include two
contradictory requirements, such as x + y :S 1 and x + y :0: 2, there can be no points
in the problem's feasible region. Linear programming problems with the empty
feasible region are called infeasible. Obviously, infeasible problems do not have
optimal solutions.
Another complication may arise if the problem's feasible region is unbounded,
as the following example demonstrates.
EXAMPLE 2 If we reverse the inequalities in problem (10.2) to x + y :0: 4 and
x + 3y :0: 6, the feasible region of the new problem will become unbounded (see
Figure 10.3). If the feasible region of a linear programming problem is unbounded,
its objective function may or may not attain a finite optimal value on it. For
example, the problem of maximizing z = 3x + Sy subject to the constraints x + y :0:
4, x + 3y :0: 6, x :0: 0, y :0: 0 has no optimal solution, because there are points in
the feasible region making 3x + Sy as large as we wish. Such problems are called
unbounded. On the other hand, the problem of minimizing z = 3x + 5y subject to
the same constraints has an optimal solution (which?).
Ill
Iterative Improvement
340
v
(0,0)
X
3x+5v=24
3x+5y=20
3x+5v=14
FIGURE 10.3 Unbounded feasible region of a linear programming problem with
constraints x + y :::: 4, x + 3y 2: 6, x 2:: 0, y 2: 0 and three level lines of
the function 3x + Sy
Fortunately, the most important features of the examples we considered above
hold for problems with more than two variables. In particular, a feasible region of
a typical linear programming problem is in many ways similar to convex polygons
in the two-dimensional Cartesian plane. Specifically, it always has a finite number
of vertices, which mathematicians prefer to call extreme points (see Section 3.3).
Furthermore, an optimal solution to a linear programming problem can be found
at one of the extreme points of its feasible region. We reiterate these properties
in the following theorem.
Any linear programming problem with
a nonempty bounded feasible region has an optimal solution; moreover, an optimal solution can always be found at an extreme point of the problem's feasible
region.2
THEOREM (Extreme Point Theorem}
2.
Except for some degenerate instances (such as maximizing z = x + y subject to x + y = 1), if a linear
programming problem with an unbounded feasible region has an optimal solution, it can also be found
at an extreme point of the feasible region.
10.1
The Simplex Method
341
This theorem implies that to solve a linear programming problem, at least
in the case of a bounded feasible region, we can ignore all but a finite number of
points in its feasible region. In principle, we can solve such a problem by computing
the value of the objective function at each extreme point and selecting the one with
the best value. There are two major obstacles to implementing this plan, however.
The first lies in the need for a mechanism for generating the extreme points of the
feasible region. As we are going to see below, a rather straightforward algebraic
procedure for this task has been discovered. The second obstacle lies in the number
of extreme points a typical feasible region has. Here, the news is bad: the number
of extreme points is known to grow exponentially with the size of the problem.
This makes the exhaustive inspection of extreme points unrealistic for most linear
programming problems of nontrivial sizes.
Fortunately, it turns out that there exists an algorithm that typically inspects
only a small fraction of the extreme points of the feasible region before reaching an
optimal one. This famous algorithm is called the simplex method. The idea of this
algorithm can be described in geometric terms as follows. Start by identifying an
extreme point of the feasible region. Then check whether one can get an improved
value of the objective function by going to an adjacent extreme point. If it is notthe
case, the current point is optimal-stop; if it is the case, proceed to an adjacent
extreme point with an improved value of the objective function. After a finite
number of steps, the algorithm will either reach an extreme point where an optimal
solution occurs or determine that no optimal solution exists.
An Outline of the Simplex Method
Our task now is to "translate" the geometric description of the simplex method
into the more algorithmically precise language of algebra. To begin with, before
we can apply the simplex method to a linear programming problem, it has to be
represented in a special form called the standardfonn. The standard form has the
following requirements:
"
It must be a maximization problem.
"
All the constraints (except the nonnegativity constraints) must be in the form
of linear equations.
"
All the variables must be required to be nonnegative.
Thus, the general linear programming problem in standard form with m constraints and n unknowns (n ::: m) is
maximize c1x1 + · · · + C11 Xn
subject to ailxl + · · · + ai 11 X11 =hi
x1 ?: 0, ... , X11 ?: 0.
[or i = 1, 2, ... , m
(10.3)
342
Iterative Improvement
It can also be written in compact matrix notations:
maximize ex
subject to Ax = b
X?::.. 0,
where
a,;,]
:
bll
[L
b
a1n
, b=
2
.
Any linear programming problem can be transformed into an equivalent
problem in standard form. If an objective function needs to be minimized, it can
be replaced by the equivalent problem of maximizing the same objective function
with all its coefficients cj replaced by -cj, j = 1, 2, ... , n (see Section 6.6 for
a more general discussion of such transformations). If a constraint is given as an
inequality, it can be replaced by an equivalent equation by adding a slack variable
representing the difference between the two sides of the original inequality. For
example, the two inequalities of problem (10.2) can be transformed, respectively,
into the following equations:
x + y + u = 4 where u ": 0 and x + 3y + v = 6 where v ": 0.
Finally, in most linear programming problems, the variables are required to be
nonnegative to begin with because they represent some physical quantities. If this
is not the case in an initial statement of a problem, an unconstrained variable
x j can be replaced by the difference between two new nonnegative variables:
"O
:::_.
xj=x1I -x 1"
,x'1O
?:. ,x
1
Thus, problem (10.2) in standard form is the following linear programming
problem in four variables:
maximize 3x + Sy + Ou + Ov
subject to x + y + u
= 4
x +3y+
+ v=6
(1o.4)
x,y,u,v2':0.
It is easy to see that if we find an optimal solution (x*, y*, u*, v*) to problem (10.4),
we can obtain an optimal solution to problem (10.2) by simply ignoring its last two
coordinates.
The principal advantage of the standard form lies in the simple mechanism
it provides for identifying extreme points of the feasible region. To do this for
problem (10.4), for example, we need to set two of the four variables in the
constraint equations to zero to get a system of two linear equations in two unknowns and solve this system. For the general case of a problem with m equations
inn unknowns (n ": m), n- m variables need to be set to zero to get a system
10.1 The Simplex Method
343
of m equations in m unknowns. If the system obtained has a unique solutionas any nondegenerate system of linear equations with the number of equations
equal to the number of unknowns does-we have a basic solation; its coordinates
set to zero before solving the system are called nonbasic, while its coordinates
obtained by solving the system are called basic. (This terminology comes from
linear algebra. Specifically, we can rewrite the system of constraint equations of
(10.4) as
A basis in the two-dimensional vector space is composed of any two vectors that
are not proportional to each other; once a basis is chosen, any vector can be
uniquely expressed as a sum of multiples of the basis vectors. Basic and nonbasic variables indicate which of the given vectors are, respectively, included and
excluded in a particular basis choice.)
If all the coordinates of a basic solution are nonnegative, the basic solution is
called a basic feasible solution. For example, if we set to zero variables x and y
and solve the resulting system for u and v, we obtain the basic feasible solution
(0, 0, 4, 6); if we set to zero variables x and u and solve the resulting system for y
and v, we obtain the basic solution (0, 4, 0, -6), which is not feasible. The importance of basic feasible solutions lies in the one-to-one correspondence between
them and the extreme points of the feasible region. For example, (0, 0, 4, 6) is an
extreme point of the feasible region of problem (10.4) (with the point (0, 0) of Figure 10.1 being its projection on the x, y plane). Incidentally, (0, 0, 4, 6) is a natural
starting point for the simplex method's application to this problem.
As mentioned above, the simplex method progresses through a series of
adjacent extreme points (basic feasible solutions) with increasing values of the
objective function. Each such point can be represented by a simplex tableau, a
table storing the information about the basic feasible solution corresponding to
the extreme point. For example, the simplex tableau for (0, 0, 4, 6) of problem
(10.4) is presented below:
X
y
"
v
I
o I
u
4
I
I
I
3
II 6
o
I
---------------------+----
-3
-s
o
o
1
1
I
t
o
(10.5)
r
---,
!
H~i' I
·.1.
,;- ,,
!·
.•.·••·••
••. '•.•.
••••••
ii'!
344
Iterative Improvement
•.••.
I
In general, a simplex tableau for a linear programming problem in standard form
with n unknowns and m linear equality constraints (n ?: m) has m + 1 rows and
n + 1 columns. Each of the first m rows of the table contains the coefficients of
a corresponding constraint equation, with the last column's entry containing the
equation's right-hand side. The columns, except the last one, are labeled by the
names of the variables. The rows are labeled by the basic variables of the basic
feasible solution the tableau represents; the values of the basic variables of this
solution are in the last column. Also note that the columns labeled by the basic
variables form the m-by-m identity matrix.
The last row of a simplex tableau is called the objective row. It is initialized
by the coefficients of the objective function with their signs reversed (in the first
n columns) and the value of the objective function at the initial point (in the last
column). On subsequent iterations, the objective row is transformed the same
way as all the other rows. The objective row is used by the simplex method to
check whether the current tableau represents an optimal solution: it does if all
the entries in the objective row-except, possibly, the one in the last column-are
nonnegative. If this is not the case, any of the negative entries indicates a nonbasic
variable that can become basic in the next tableau.
For example, according to this criterion, the basic feasible solution (0, 0, 4, 6)
represented by tableau (10.5) is not optimal. The negative value in the x-column
signals the fact that we can increase the value of the objective function z = 3x +
Sy + Ou + Ov by increasing the value of the x-coordinate in the current basic
feasible solution (0, 0, 4, 6). Indeed, since the coefficient for x in the objective
function is positive, the larger the x value, the larger the value of this function. Of
course, we will need to "compensate" an increase in x by adjusting the values of
the basic variables u and v so that the new point is still feasible. For this to be the
case, both following conditions must be satisfied:
x +u = 4
where u ?: 0
x + v = 6 where v ?: 0,
which means that
x s min{4, 6] = 4.
Note that if we increase the value of x from 0 to 4, the largest amount possible,
we will find ourselves at the point (4, 0, 0, 2), an adjacent to (0, 0, 4, 6) extreme
point of the feasible region, with z = 12.
Similarly, the negative value in the y-column of the objective row signals the
fact that we can also increase the value of the objective function by increasing
the value of they-coordinate in the initial basic feasible solution (0, 0, 4, 6). This
reqnires the following:
y + u = 4 where u ?: 0
3y + v = 6 where v ?: 0,
10.1
The Simplex Method
345
which means that
. {4 6}
.y <
- mm -.,
1 -3 = 2 .
If we increase the value of y from 0 to 2, the largest amount possible, we will find
ourselves at the point (0, 2, 2, 0), another adjacent to (0, 0, 4. 6) extreme point,
with z = 10.
If there are several negative entries in the objective row, a commonly used
rule is to select the most negative one, i.e., the negative number with the largest
absolute value. This rule is motivated by the observation that such a choice yields
the largest increase in the objective function's value per unit of change in a variable's value. (In our example, an increase in the x-value from 0 to 1 at (0, 0, 4, 6)
changes the value of z = 3x + 5y + Ou + Ov from 0 to 3, while an increase in the
y-value from 0 to 1 at (0. 0, 4, 6) changes z from 0 to 5.) Note, however, that the
feasibility constraints impose different limits on how much each of the variables
may increase. In our example, in particular, the choice of the y-variable over the
x-variable leads to a smaller increase in the value of the objective function. Still,
we will employ this commonly used rule and select variable y as we continue with
our example. A new basic variable is called the entering variable, while its column
is referred to as the pivot column; we mark the pivot column by t .
Now we will explain how to choose a departing variable, i.e., a basic variable
to become non basic in the next tableau. (The total number of basic variables in any
basic solution must be equal to m, the number of the equality constraints.) As we
saw above, to get to an adjacent extreme point with a larger value of the objective
function, we need to increase the entering variable by the largest amount possible
to make one of the old basic variables zero while preserving the nonnegativity
of all the others. We can translate this observation into the following rule for
choosing a departing variable in a simplex tableau: For each positive entry in the
pivot column, compute the &-ratio by dividing the row's last entry by the entry in
the pivot column. For the example of tableau (10.5), these &-ratios are:
e, = 4 =4, e, =-6 =2.
1
3
The row with the smallest 8-ratio determines the departing variable, i.e., the
variable to become nonbasic. Ties may be broken arbitrarily. For our example, it is
variable v. We mark the row ofthe departing variable, called the pivot row, by<-and denote it ffiw. Note that if there are no positive entries in the pivot column,
no 8-ratio can be computed, which indicates that the problem is unbounded and
the algorithm stops.
Finally, the following steps need to be taken to transform a current tableau
into the next one. (This transformation, called pivoting, is similar to the principal step of the Gauss-Jordan elimination algorithm for solving systems of linear
equations-see Problem 8 in Exercises 6.2.) First, divide all the entries of the pivot
'+
346
Iterative Improvement
row
row by the pivot, its entry in the pivot column, to obtain
new· For tableau
(10.5), we obtain
1
1
1 0
2.
3
3
Then replace each of the other rows, including the objective row, by the
difference
where cis the row's entry in the pivot column. For tableau (10.5), this yields
2
row 1 - 1 . rownew:
row 3- (-5).
row
3
4
new=
3
0
1
0
0
1
3
5
3
--
2,
10.
Thus, the simplex method transforms tableau (10.5) into the following tableau:
X
)'
2
0
u
v
_l_
3
3
I
I
2
:
I
I
)'
1
0
3
1
3
:
I
2
(10.6)
I
---------------------+---4
-3
0
0
Q
3
1
:
10
t
Tableau (10.6) represents the basic feasible solution (0, 2, 2, 0) with an increased
value of the objective function, which is equal to 10. It is not optimal, however
(why?).
The next iteration-do it yourself as a good exercise!-yields tableau (1 0.7):
X
y
u
0
X
v
I
I
_.1_
2
:
.l
I
1
3
I
)'
-,
1
0
2
(10.7)
I
I
---------------------+---1
0
0
2
:
I
14
10.1
The Simplex Method
347
This tableau represents the basic feasible solution (3, 1, 0, 0). It is optimal because
all the entries in the objective row of tableau (10.7) are nonnegative. The maximal
value of the objective function is equal to 14, the last entry in the objective row.
Let us summarize the steps of the simplex method.
Summary of the simplex method
Step 0 Initialization Present a given linear programming problem in standard form and set up an initial tableau with nonnegative entries in the
rightmost column and m other columns composing the m-by-m identity
matrix. (Entries in the objective row are to be disregarded in verifying
these requirements.) These m columns define the basic variables of the
initial basic feasible solution, used as the labels of the tableau's rows.
Step 1 Optimality test If all the entries in the objective row (except, possibly,
the one in the rightmost column, which represents the value of the
objective function) are nonnegative-stop: the tableau represents an
optimal solution whose basic variables' values are in the rightmost
column and the remaining, nonbasic variables' values are zeros.
Step 2 Finding the entering variable Select a negative entry among the first
n elements of the objective row. (A commonly used rule is to select
the negative entry with the largest absolute value, with ties broken
arbitrarily.) Mark its column to indicate the entering variable and the
pivot column.
Step 3 Finding the departing variable For each positive entry in the pivot column, calculate the &-ratio by dividing that row's entry in the rightmost
column by its entry in the pivot column. (If all the entries in the pivot
column are negative or zero, the problem is unbounded-stop.) Find
the row with the smallest &-ratio (ties may be broken arbitrarily), mark
this row to indicate the departing variable and the pivot row.
Step 4 Forming the next tableau Divide all the entries in the pivot row by
its entry in the pivot column. Subtract from each of the other rows,
including the objective row, the new pivot row multiplied by the entry
in the pivot column of the row in question. (This will make all the
entries in the pivot column O's except for 1 in the pivot row). Replace
the label of the pivot row by the variable's name of the pivot column
and go back to Step 1.
Further Notes on the Simplex Method
Formal proofs of validity of the simplex method steps can be found in books
devoted to a detailed discussion of linear programming (e.g., [Dan63]). A few
important remarks about the method still need to be made, however. Generally
speaking, an iteration of the simplex method leads to an extreme point of the problem's feasible region with a greater value of the objective function. In degenerate
In
1.
!!!{:
!
'!·1(···1
l'.
'-:<
j,·'..·'
I
!
348
Iterative Improvement
cases, which arise when one or more basic variables are equal to zero, the simplex
method can only guarantee that the value of the objective function at the new
extreme point is greater than or equal to its value at the previous point. In turn,
this opens a door to the possibility that not only the objective function's values
"stall" for several iterations in a row but that the algorithm might cycle back to a
previously considered point and hence never terminate. The latter phenomenon
is called cycling. Although it rarely if ever happens in practice, specific examples
of problems where cycling does occur have been constructed. A simple modification of Steps 2 and 3 of the simplex method, called Bland's rule, eliminates even
the theoretical possibility of cycling. Assuming that the variables are denoted by
a subscripted letter (e.g., XJ, x 2 , ... , x,), this rule can be stated as follows:
Step 2 modified Among the columns with a negative entry in the objective
row, select the column with the smallest subscript.
Step 3 modified Resolve a tie among the smallest &-ratios by selecting the
row labeled by the basic variable with the smallest subscript.
Another caveat deals with the assumptions made in Step 0. They are automatically satisfied if a problem is given in the form where all the constraints imposed
on nonnegative variables are inequalities a 11x1 + · · · + a 111 x 11 .:::: b1 with b1 ::: 0 for
i = 1, 2, ... , m. Indeed, by adding a nonnegative slack variable x,+; into the ith
constraint, we obtain the equality anx 1 + · · · + ai 11 x 11 + xn+i = b1, and all therequirements imposed on an initial tableau of the simplex method are satisfied for
the obvious basic feasible solution Xt = · · · = x 11 = 0, Xn+l = · · · = Xn+m = 1. But
if a problem is not given in such a form, finding an initial basic feasible solution
may present a nontrivial obstacle. Moreover, for problems with an empty feasible
region, no initial basic feasible solution exists, and we need an algorithmic way to
identify such problems. One of the ways to address these issues is to use an extension to the classic simplex method called the two-phase simplex method (see, e.g.,
[Kol95]). In a nutshell, this method adds a set of artificial variables to the equality
constraints of a given problem so that the new problem has an obvious basic feasible solution. It then solves the linear programming problem of minimizing the
sum of the artificial variables by the simplex method. The optimal solution to this
problem either yields an initial tableau for the original problem or indicates that
the feasible region of the original problem is empty.
How efficient is the simplex method? Since the algorithm progresses through
a sequence of adjacent points of a feasible region, one should probably expect bad
news because the number of extreme points is known to grow exponentially with
the problem size. Indeed, the worst-case efficiency of the simplex method has been
shown to be exponential as well. Fortunately, more than half a century of practical
experience with the algorithm has shown that the number of iterations in a typical
application ranges between m and 3m, with the number of operations per iteration
proportional to mn, where m and n are the numbers of equality constraints and
variables, respectively.
10.1
The Simplex Method
349
Since its discovery in 1947, the simplex method has been a subject of intensive
study by many researchers. Some of them have worked on improvements to the
original algorithm and details of its efficient implementation. As a result of these
efforts, programs implementing the simplex method have been polished to the
point that very large problems with hundreds of thousands of constraints and
variables can be solved in a routine manner. In fact, such programs have evolved
into sophisticated software packages. These packages enable the user to enter
a problem's constraints and obtain a solution in a user-friendly form. They also
provide tools for investigating important properties of the solution such as its
sensitivity to changes in the input data. Such investigations are very important for
many applications, including those in economics. At the other end of the spectrum,
linear programming problems of a moderate size can nowadays be solved on a
desktop using a standard spreadsheet facility or by taking advantage of specialized
software available on the Internet.
Researchers have also tried to fmd algorithms for solving linear programming
problems with polynomial-time efficiency in the worst case. An important milestone in the history of such algorithms was the proof by L. G. Khachian [Kha79]
showing that the ellipsoid method can solve any linear programming problem in
polynomial time. Although the ellipsoid method was much slower than the simplex method in practice, its better worst-case efficiency encouraged a search for
alternatives to the simplex method. In 1984, Narendra Karmarkar published an
algorithm that not only had a polynomial worst-case efficiency but was competitive with the simplex method in empirical tests as well. Although we are not going
to discuss Karmarkar's algorithm [Kar84] here, it is worth pointing out that it is
also based on the iterative improvement idea; however, it generates a sequence
of feasible solutions that lie within the feasible region rather than going through a
sequence of adjacent extreme points as the simplex method does. Such algorithms
are called interior-point methods (see, e.g., [Ye97]).
-----Exercises 10.1 - - - - - - - - - - - - - - 1. Solve the following linear programming problems geometrically.
a.
maximize 3x + y
subject to -x + y :=:: 1
2x+y:o4
X 2:
b.
0, y 2: 0
maximize x + 2 y
subject to 4x 2: y
y:03+x
X 2: 0, y 2: 0
350
Iterative Improvement
2. Consider the linear programming problem
minimize c 1x + c2y
subject to
x + y ": 4
X
+ 3y ": 6
X ":
0, y ": 0
where c 1 and c2 are some real numbers not both equal to zero.
a. Give an example of the coefficient values c1 and c2 for which the problem
has a unique optimal solution.
b. Give an example of the coefficient values c1 and c2 for which the problem
has infinitely many optimal solutions.
c. Give an example of the coefficient values c1 and c2 for which the problem
does not have an optimal solution.
3. Would the solution to problem (10.2) be different if its inequality constraints
were strict, i.e., x + y < 4 and x + 3y < 6, respectively?
4. Trace the simplex method on
a. the problem of Exercise 1a.
b. the problem of Exercise 1b.
5. Trace the simplex method on the problem of Example 1 in Section 6.6
a. by hand.
b. by using one of the implementations available on the Internet.
6. Determine how many iterations the simplex method needs to solve the
problem
n
maximize
Lx
j
}=1
subject to 0 <( xj <( bj, where bj > 0 for j = 1, 2, ... , n.
7. Can we apply the simplex method to solve the knapsack problem (see Example 2 in Section 6.6)? If you answer yes, indicate whether it is a good algorithm
for the problem in question; if you answer no, explain why not.
8. Prove that no linear programming problem can have exactly k :, 1 optimal
solutions unless k = 1.
10.2
The Maximum-Flow Problem
351
9. If a linear programming problem
maximize
L" c1x1
J=l
subject to
L" aijx1 .:::; hi fori = 1, 2, ... , m
)=1
x 1 , Xz, .. , , X 11 :::: 0
is considered as primal, then its dual is defmed as the linear programming
problem
minimize
L"' hiYi
i=l
subject to
L"' aijYi ::=: c1 for j = 1, 2, ... , n
i=l
Yt • Yz, · · · , Y111 ~ 0.
a. Express the primal and dual problems in matrix notations.
b. Find the dual of the following linear programming problem
maximize x1 + 4xz -
x3
subject to x1 + x 2 + x 3 s 6
x1 - x2 - lx3 ::; 2
x 1, Xz, XJ ~ 0.
c. Solve the primal and dual problems and compare the optimal values of
their objective functions.
In a parliament, each parliamentarian has at most
three enemies. Design an algorithm that divides the parliament into two
10. Parliament pacification
chambers in such a way that no parliamentarian has more than one enemy
in his or her chamber. (after [Sav03], p. 1, #4)
10.2 The Maximum-Flow Problem
In this section, we consider the important problem of maximizing the flow of
a material through a transportation network (pipeline system, communication
system, and so on). We will assume that the transportation network in question
can be represented by a connected weighted digraph with n vertices numbered
from 1 to n and a set of edges E, with the following properties:
Iterative Improvement
352
5
3
4
2
2
)-----(6
3
4
FIGURE 10.4 Example of a network graph. The vertex numbers are vertex "names;''
the edge numbers are edge capacities.
"
It contains exactly one vertex with no entering edges; this vertex is called the
"
It contains exactly one vertex with no leaving edges; this vertex is called the
sink and assumed to be numbered n .
The weight uij of each directed edge (i, j) is a positive integer, called the
source and assumed to be numbered 1.
.,
edge capacity. (This number represents the upper bound on the amount of
the material that can be sent from i to j through a link represented by this
edge).
A digraph satisfying these properties is called a flow network or simply a
network 3 A small instance of a network is given in Figure 10.4.
It is assumed that the source and the sink are the only source and destination
of the material, respectively; all the other vertices can serve only as points where
a flow can be redirected without consuming or adding any amount of the material.
In other words, the total amount of the material entering an intermediate vertex
must be equal to the total amount of the material leaving the vertex. This condition is called the flow-conservation requirement. If we denote the amount sent
through edge (i, j) by xij, then for any intermediate vertex i, the flow-conservation
requirement can be expressed by the following equality constraint:
L x1
;=
j: (j,i)EE
L xij fori=2,3, ... ,n-l,
(10.8)
i: (i,j)EE
where the sums in the left- and right-hand sides express the total inflow and outflow
entering and leaving vertex i, respectively.
Since no amount of the material can change by going through intermediate
vertices of the network, it stands to reason that the total amount of the material
leaving the source must end up at the sink. (This observation can also be derived
3.
In a slightly more general model, one can consider a network with several sources and sinks and allow
capacities uiJ to be infinitely large.
10.2
The Maximum-Flow Problem
353
formally from equalities (10.8), a task you will be asked to do in the exercises.)
Thus, we have the following equality:
I:
Xtj
=
j:(l,j)EE
I:
X jw
(10.9)
j:(j,n)EE
This quantity, the total outflow from the source-or, equivalently, the total inflow
into the sink-is called the value of the flow. We denote it by v. It is this quantity
that we will want to maximize over all possible flows in a network.
Thus, a (feasible) flow is an assignment of real numbers xu to edges (i, j) of
a given network that satisfy flow-conservation constraints (10.8) and the capacity
constraints
0 :<= x 0 :<= u 0
for every edge (i, j) E E.
(10.10)
The maximum-flow problem can be stated formally as the following optimization
problem:
maximize v =
L x1 1
j: (1 ,))EE
subject to
L x.i
i
L xu = 0 for i = 2, 3, ... , n - 1 (10.11)
): (j,i)EE
j: (i,j)EE
Osxij:<=uij
foreveryedge(i,j)EE.
We can solve linear programming problem (10.11) by the simplex method or
by another algorithm for general linear programming problems (see Section 10.1 ).
However, the special structure of problem (10.11) can be exploited to design faster
algorithms. In particular, it is quite natural to employ the iterative improvement
idea. We can always start with the zero flow (i.e., set xu= 0 for every edge (i, j)
in the network). Then, on each iteration, we can try to find a path from source
to sink along which some additional flow can be sent. Such a path is called flow
augmenting. If a flow-augmenting path is found, we adjust the flow along the edges
of this path to get a flow of an increased value and try to find an augmenting
path for the new flow. If no flow-augmenting path can be found, we conclude
that the current flow is optimal. This general template for solving the maximumflow problem is called the augmenting-path method, also known as the FordFulkerson method after L. R. Ford, Jr. and D. R. Fulkerson, who discovered it
(see [For57]).
An actual implementation of the augmenting path idea is, however, not quite
straightforward. To see this, let us consider the network in Figure 10.4. We start
with the zero flow shown in Figure 10.5a. (In that figure, the zero amounts sent
through each edge are separated from the edge capacities by the slashes; we will
use this notation in the other examples as well.) It is natural to search for a flowaugmenting path from source to sink by following directed edges (i, j) for which
the current flow xij is less than the edge capacity uu. Among several possibilities,
let us assume that we identify the augmenting path 1-+2-+3-+6 first. We can
m
I'
i
354
Iterative Improvement
I'
I'"
[ r
r:
I
5
0/3
0/4
0/2
0/2
l-__,;=--t 6
0/3
0/1
4
(a)
(b)
5
1/3
1/4
2/2
2/2
1/3
6
1/1
4
(c)
FIGURE 10.5 Illustration of the augmenting-path method. Flow-augmenting paths are
shown in bold. The flow amounts and edge capacities are indicated by
the numbers before and after the slash, respectively.
increase the flow along this path by a maximum of 2 units, which is the smallest
unused capacity of its edges. The new flow is shown in Figure 10.5h. This is as far
as our simpleminded idea about flow-augmenting paths will be able to take us.
Unfortunately, the flow shown in Figure 10.5b is not optimal: its value can still be
increased along the path 1-+4-+3+-2-+5-+6 by increasing the flow by 1 on edges
355
10.2 The Maximum-Flow Problem
(1.4), (4,3), (2,5), and (5,6) and decreasing it by 1 on edge (2,3). The flow obtained
as the result of this augmentation is shown in Figure 10.5c. It is indeed maximal.
(Can you tell why?)
Thus, to find a flow-augmenting path for a flow x, we need to consider paths
from source to sink in the underlying undirected graph in which any two consecutive vertices i, j are either
(i) connected by a directed edge from i to j with some positive unused capacity
riJ = u;j -xi} (so that we can increase the flow through that edge by up to riJ units)
or
(ii) connected by a directed edge from j to i with some positive flow xji (so that
we can decrease the flow through that edge by up to xji units).
Edges of the first kind are called forward edges because their tail is listed before
their head in the vertex list 1 -> ... i -> j ... -+ n defining the path; edges of the
second kind are called backward edges because their tail is listed after their head
in the path list 1 -+ ... i <-- j ... -> n. To illustrate, for the path 1 -> 4 -> 3 <-- 2 ->
5-> 6 of the last example, (1,4), (4,3), (2,5), and (5,6) are the forward edges, while
(3,2) is the backward edge.
For a given flow-augmenting path, let r be the minimum of all the unused
capacities r;j of its forward edges and all the flows xji of its backward edges.
It is easy to see that if we increase the current flow by r on each forward edge
and decrease it by this amount on each backward edge, we will obtain a feasible
flow whose value is r units greater than the value of its predecessor. Indeed, let
i be an intermediate vertex on a flow-augmenting path. There are four possible
combinations offorward and backward edges incident to vertex i:
+r . +r
-----+ l -----+ '
+r . -r
-----+ l +----,
-r . +r
+ - l -----+,
-r
.
-r
+-t+--.
For each of them, the flow-conservation requirement for vertex i will still hold
after the flow adjustments indicated above the edge arrows. Further, since r is the
minimum among all the positive unused capacities on the forward edges and all
the positive flows on the backward edges of the flow-augmenting path, the new
flow will satisfy the capacity constraints as well. Finally, adding r to the flow on
the first edge of the augmenting path will increase the value of the flow by r.
Under the assumption that all the edge capacities are positive integers, r
will be a positive integer too. Hence, the flow value increases at least by 1 on
each iteration of the augmenting-path method. Since the value of a maximum
flow is bounded above (e.g., by the sum of the capacities of the edges incident
to the source), the augmenting-path method has to stop after a finite number of
iterations. 4 Surprisingly, the final flow always turns out to be maximal, irrespective
of a sequence of augmenting paths. This remarkable result stems from the proof
4.
If capacity upper bounds are irrational numbers, the augmenting-path method may not terminate
(see, e.g., [Chv83], pp. 387-388, for a cleverly devised example demonstrating such a situation). 'I11is
356
Iterative Improvement
0/U
1/U
0/U
0/U
4
0/U
0/U
1/U
1/U
4
1/U
0/U
(a)
4
1/U
(b)
U/U
1/U
(c)
U/U
4
U/U
U/U
(d)
FIGURE 10.6 Efficiency degradation of the augmenting-path method
of the Max-Flow Min-Cut Theorem (see, e.g., [For62)), which we replicate later
in this section.
The augmenting-path method-as described above in its general form-does
not indicate a specific way for generating flow-augmenting paths. A bad sequence
of such paths may, however, have a dramatic impact on the method's efficiency.
Consider, for example, the network in Figure 10.6a, in which U stands for some
large positive integer. If we augment the zero flow along the path 1--+2--+3--+4,
we shall obtain the flow of value l shown in Figure 10.6b. Augmenting that flow
along the path 1--+3<--2--+4 will increase the flow value to 2 (Figure 10.6c). If
we continue selecting this pair of flow-augmenting paths, we will need a total of
2U iterations to reach the maximum flow of value 2U (Figure 10.6d). Of course,
we can obtain the maximum flow in just two iterations by augmenting the initial
zero flow along the path 1--+ 2--+4 followed by augmenting the new flow along the
path 1--+3--+4. The dramatic difference between 2U and 2 iterations makes the
point.
Fortunately, there are several ways to generate flow-augmenting paths efficiently and avoid the degradation in performance illustrated by the previous
limitation is only of theoretical interest because we cannot store irrational numbers in a computer, and
rational numbers can be transformed into integers by changing the capacity measurement unit.
10.2
The Maximum·Fiow Problem
357
example. The simplest of them uses breadth-first search to generate augmenting paths with the least number of edges (see Section 5.2). This version of the
augmenting-path method, called shortest-augmenting-path or first-labeled-first·
scanned algorithm, was suggested by J. Edmonds and R. M. Karp [Edm72]. The
labeling refers to marking a new (unlabeled) vertex with two labels. The first label
indicates the amount of additional flow that can be brought from the source to
the vertex being labeled. The second label is the name of the vertex from which
the vertex being labeled was reached. (It can be left undefined for the source.) It
is also convenient to add the +or - sign to the second label to indicate whether
the vertex was reached via a forward or backward edge, respectively. The source
can be always labeled with oo, -.For the other vertices, the lables are computed
as follows.
If unlabeled vertex j is connected to the front vertex i of the traversal queue
by a directed edge from i to j with positive unused capacity rij = uij- xij, then
vertex j is labeled with lj, ;+,where l; =min {I;, rij).
If unlabeled vertex j is connected to the front vertex i of the traversal queue
by a directed edge from j to i with positive flow xji• then vertex j is labeled with
lj, ;-,where lj =min {I;, xj;).
If this labeling-enhanced traversal ends up labeling the sink, the current
flow can be augmented by the amount indicated by the sink's first label. 'The
augmentation is performed along the augmenting path traced by following the
vertex second labels from sink to source; the current flow quantities are increased
on the forward edges and decreased on the backward edges of this path. If, on the
other hand, the sink remains unlabeled after the traversal queue becomes empty,
the algorithm returns the current flow as maximal and stops.
ALGORITHM
ShortestAugmentingPath(G)
//Implements the shortest-augmenting-path algorithm
//Input: A network G with single source 1, single sink n, and
positive integer capacities uij on its edges (i, j)
//Output: A maximum flow x
assign xij = 0 to every edge (i, j) in the network
label the source with oo, - and add the source to the empty queue Q
while not Empty(Q) do
i <- Front(Q); Dequeue(Q)
for every edge from i to j do //forward edges
if j is unlabeled
rij +-- Uij -
Xij
if ,.ij > 0
lj <-min {I;, rij); label j with lj, ;+
Enqueue(Q, j)
r
~.
B1!
1
•
.•-._:; ._:._ !.• 1.
A:r ~
358
Iterative Improvement
for every edge from j to i do //backward edges
if j is unlabeled
ifxji>O
11 <-min[!,, xfi); label j with t1, iEnqueue(Q, j)
if the sink has been labeled
//augment along the augmenting path found
j <- n //start at the sink aud move backwards using second labels
while j i" 1 //the source hasn't been reached
if the second label of vertex j is i+
Xij +--- Xij
+l
11
else //the second label of vertex j is ixji +-- xji - l 11
j <- i; i <-the vertex indicated by i's second label
erase all vertex labels except the ones of the source
reinitialize Q with the source
return x //the current flow is maximum
An application of this algorithm to the network in Figure 10.4 is illustrated in
Figure 10.7.
The optimality of a fmal flow obtained by the augmenting-path method stems
from a theorem that relates network flows to network cuts. A cut induced by
partitioning vertices of a network into some subset X containing the source and
X, the complement of X, containing the sink is the set of all the edges with a tail
in X and a head in X. We denote a cut C(X, X) or simply C. For example, for the
network in Figure 10.4, each of the following subsets is the cut
C(X, X)= {(1, 2), (1, 4)] if X= (1] and hence X= {2, 3, 4, 5, 6},
C(X, X)= {(3, 6), (5, 6)] if X= {1, 2, 3, 4, 5] and hence X= {6],
C(X, X)= {(2, 5), (2, 3), (4, 3)} if X= {1, 2, 4] andhencei· = {3, 5, 6].
The name "cut" stems from the following property: if all the edges of a cut
were deleted from the network, there would be no directed patb from source to
sink. Indeed, let C (X, X) be a cut. Consider a directed path from source to sink.
If is the first vertex of that patb which belongs to X (the set of such vertices is
not empty because it contains the sink), then its immediate predecessor v,_ 1 on
that path belongs to X, and hence the edge from
must be an element of
1 to
the cut C (X, X). This proves the property in question.
The capacity of a cut C(X, X), denoted c(X, X), is defined as the sum of
capacities of the edges that compose the cut. For the three examples of cuts given
above, the capacities are equal to 5, 6, and 9, respectively. Since the number of
different cuts in a network is nonempty and finite (why?), there always exists
a minimum cut, i.e., a cut with the smallest capacity. (What is a minimum cut
v,
v, _ v,
i
~
1
r·
1
10.2
The Maximum-Flow Problem
2, 2+
5
5
0/3
0/3
359
0/4
0/3
0/1
0/4
0/1
4
4
3, 1+
Queue:l24356
tt t t
Augment the flow by 2 (the sink's first label)
along the path I _,. 2 -7 3 -7 6.
5
0/3
0/3
0/4
0/1
4
3, 1+
Queue:l43256
t t t t t
Augment the flow by I (the sink's first label)
along the path 1 ---7 4 -7 3 f-. 2 -----7 5 -7 6.
5
1/3
1/3
5
1/4
1/1
4
Queue: 1 4
t t
1/3
1/3
1/4
1/1
4
2, 1+
No augmenting path (the sink is unlabeled);
the current flow is maximal.
FIGURE 10.7 Illustration of the shortest-augmenting-path algorithm. The diagrams on
the left show the current flow before the next iteration begins; the
diagrams on the right show the results of the vertex labeling on that
iteration, the augmenting path found lin bold), and the flow before its
augmentation. Vertices deleted from the queue are indicated by the t
sign.
360
Iterative Improvement
in the network of Figure 10.4?) The following theorem establishes an important
relationship between the notions of maximum flow and minimum cut
THEOREM (Max-Flow Min-Cut Theorem) The value of a maximum flow in a
network is equal to the capacity of its minimum cut.
PROOF First, let x be a feasible flow of value v and let C(X, X) be a cut of
capacity c in the same network. Consider the flow across this cut defined as the
difference between the sum of the flows on the edges from X to X and the snm
of the flows on the edges from X to X. It is intuitively clear and can be formally
derived from the equations expressing the flow-conservation requirement and the
definition of the flow value (Problem 6(b) in the exercises) that the flow across the
cut C(X, X) is equal to v, the value of the flow:
(10.12)
iEX, }EX
Since the second sum is nonnegative and the flow xiJ on any edge (i, j) cannot
exceed the edge capacity uiJ, equality (10.12) implies that
iEX, }Ei
i.e.,
v.::: c.
(10.13)
Thus, the value of any feasible flow in a network cannot exceed the capacity of
any cut in that network.
Let v* be the value of a final flow x* obtained by the augmenting-path method.
If we now find a cut whose capacity is equal to v*, we will have to conclude, in view
of inequality (10.13), that (i) the value v' of the final flow is maximal among all
feasible flows; (ii) the cut's capacity is minimal among all cuts in the network;
(iii) the maximum-flow value is equal to the minimum-cut capacity.
To find such a cut, consider the set of vertices X* that can be reached from the
source by following an undirected path composed of forward edges with positive
unused capacities (with respect to the final flow x') and backward edges with
positive flows on them. This set contains the source but does not contain the
sink: if it did, we would have an augmenting path for the flow x*, which would
contradict the assumption that the flow x* is final. Consider the cut CCX*, X'). By
the definition of set X*, each edge (i, j) from X* to X* has zero unused capacity,
i.e., x;j = uiJ, and each edge (j, i) from X' to X' has the zero flow on it (otherwise,
j would be in X*). Applying equality (10.12) to the final flow x' and the set X*
defined above, we obtain
10.2
The Maximum-Flow Problem
v* =
361
c(X*, X*),
)EX*, iEX*
which proves the theorem.
T11e proof outlined above accomplishes more than proving the equality of
the maximum-flow value and the minimum-cut capacity. It also implies that when
the augmenting-path method terminates, it yields both a maximum flow and a
minimum cut. If labeling of the kind utilized in the shortest-augmenting-path
algorithm is used, a minimum cut is formed by the edges from the labeled to
unlabeled vertices on the last iteration of the method. Finally, the proof implies
that all such edges must be full (i.e., the flows must be equal to the edge capacities),
while all the edges from unlabeled vertices to labeled, if any, must be empty
(i.e., have zero flows on them). In particular, for the network of Figure 10.7, the
algorithm finds the cut {(1, 2), (4, 3)) of minimum capacity 3, both edges of which
are full as required.
Edmonds and Karp proved in their paper [Edm72] that the number of augmenting paths needed by the shortest-augmenting-path algorithm never exceeds
nm /2, where n and m are the number of vertices and edges, respectively. Since
the time required to find a shortest augmenting path by breadth-first search is
in O(n + m) = O(m) for networks represented by their adjacency lists, the time
efficiency of the shortest-augmenting-path algorithm is in O(nm 2 ).
More efficient algorithms for the maximum-flow problem are known (see the
monograph [Ahu93] as well as appropriate chapters in such books as [CarOl]
and [Sed98]). Some of them implement the augmenting-path idea in a more efficient manner. Others are based on the concept of preftows. A preflow is a flow
that satisfies the capacity constraints but not the flow-conservation requirement.
Any vertex is allowed to have more flow entering the vertex than leaving it. A
preflow-push algorithm moves the excess flow toward the sink until the flowconservation requirement is re-established for all intermediate vertices of the
network. Faster algorithms of this kind have worst-case efficiency close to O(nm).
Note that preftow-push algorithms fall outside the iterative-improvement paradigm because they do not generate a sequence of improving solutions that satisfy
all the constraints of the problem.
To conclude this section, it is worth pointing out that although the initial
interest in studying network flows was caused by transportation applications, this
model has also proved to be useful for many other areas. We discuss one of iliem
in the next section.
-----Exercises 1 1 0 . 2 - - - - - - - - - - - - - - 1, Since maximum-flow algorithms require processing edges in both directions,
it is convenient to modify the adjacency matrix representation of a network
I
362
Iterative Improvement
as follows. If there is a directed edge from vertex ito vertex j of capacity uij,
then the element in the ith row and the jth column is set to uij, while the
element in the jth row and the ith column is set to -uij; if there is no edge
between vertices i and j, both these elements are set to zero. Outline a simple
algorithm for identifying a source and a sink in a network presented by such
a matrix and indicate its time efficiency.
2. Apply the shortest-augmenting-path algorithm to find a maximum flow and a
minimum cut in the following networks.
a.
5
2
2
4
6
3
4
4
7
5
8
6
b.
2
4
7
6
4
5
3. a. Does the maximum-flow problem always have a unique solution? Would
your answer be different for networks with different capacities on all their
edges?
b. Answer the same questions for the minimum-cut problem of finding a cut
of the smallest capacity in a given network.
4. a. Explain how the maximum-flow problem for a network with several
sources and sinks can be transformed into the same problem for a network
with a single source and a single sink.
b. Some networks have capacity constraints on the flow amounts that can
flow through their intermediate vertices. Explain how the maximum-flow
problem for such a network can be transformed into the maximum-flow
problem for a network with edge capacity constraints only.
5. Consider a network that is a rooted tree, with the root as its source, the leaves
as its sinks, and all the edges directed along the paths from the root to the
leaves. Design an efficient algorithm for finding of a maximum flow in such a
network. What is the time efficiency of your algorithm?
6. a. Prove equality (10.9).
b. Prove that for any flow in a network and any cut in it, the value of the
flow is equal to the flow across the cut (see equality (10.12)). Explain the
relationship between this property and equality (10.9).
I
10.3
Maximum Matching in Bipartite Graphs
363
7. a. Express the maximum-flow problem for the network of Figure 10.4 as a
linear programnting problem.
b. Solve this linear programming problem by the simplex method.
8. As an alternative to the shortest-augmenting-path algorithm, Edmonds and
Karp [Edm72] suggested the maximum-capacity-augmenting-path algorithm,
in which a flow is augmented along the path that increases the flow by the
largest amount. Implement both these algorithms in the language of your
choice and perform an empirical investigation of their relative efficiency.
9. Write a report on a more advanced maximum-flow algorithm such as (i)
Dinitz's algorithm; (ii) Karzanov's algorithm; (iii) Malhotra-KamarMaheshwari algorithm; or (iv) Goldberg-Tarjan algorithm.
10. Dining problem Several families go out to dinner together. To increase their
social interaction, they would like to sit at tables so that no two members of
the same family are at the same table. Show how to find a seating arrangement
that meets this objective (or prove that no such arrangement exists) by using
a maximum-flow problem. Assume that the dinner contingent has p families
and that the ith family has members. Also assume that q tables are available
and the jth table has a seating capacity of bj. [Ahu93], p. 198
a,
10.3 Maximum Matching in Bipartite Graphs
In many situations we are faced with a problem of pairing elements of two sets.
The traditional example is boys and girls for a dance, but you can easily think
of more serious applications. It is convenient to represent elements of two given
sets by vertices of a graph, with edges between vertices that can be paired. A
matching in a graph is a subset of its edges with the property that no two edges
share a vertex. A maximum matching~more precisely, a maximum cardinality
matching-is a matching with the largest number of edges. (What is it for the graph
in Figure 10.8? Is it unique?) The maximum-matching problem is the problem of
finding a maximum matching in a given graph. For an arbitrary graph, this is a
rather difficult problem. It was solved in 1965 by Jack Edmonds [Edm65]. (See
[Gal86] for a good survey and more recent references.)
We limit our discussion in this section to the simpler case of bipartite graphs. In
a bipartite graph, all the vertices can be partitioned into two disjoint sets V and U,
not necessarily of the same size, so that every edge connects a vertex in one of these
sets to a vertex in the other set. In other words, a graph is bipartite if its vertices
can be colored in two colors so that every edge has its vertices colored in different
colors; such graphs are also said to be 2-colorable. The graph in Figure 10.8 is
bipartite. It is not difficult to prove that a graph is bipartite if and only if it does
not have a cycle of an odd length. We will assume for the rest of this section that
the vertex set of a given bipartite graph has been already partitioned into sets V
and U as required by the definition (see Problem 8 in Exercises 5.2).
364
Iterative Improvement
v
4
u
8
FIGURE 10.8 Example of a bipartite graph
Let us try to apply the iterative-improvement technique to the maximumcardinality-matching problem. Let M be a matching in a bipartite graph G =
(V, U, E). How can we improve it, i.e., find a new matching with more edges?
Obviously, if every vertex in either V or U is matched (has a mate), i.e., serves as
an endpoint of an edge in M, this cannot be done and M is a maximum matching.
Therefore, to have a chance at improving the current matching, both V and U must
contain unmatched (also called free) vertices, i.e., vertices that are not incident to
any edge in M. For example, for the matching Ma = {(4, 8), (5, 9)) in the graph of
Figure 10.9a, vertices 1, 2, 3, 6, 7, and 10 are free, while vertices 4, 5, 8, and 9 are
matched.
Another obvious observation is that we can immediately increase a current
matching by adding an edge between two free vertices. For example, adding (1, 6)
to the matching M 0 = {(4, 8), (5, 9)) in the graph in Figure 10.9a yields a larger
matching Mb = {(1, 6), (4, 8), (5, 9)) (Figure 10.9b). Let us now try to find a
matching larger than Mh by matching vertex 2. The only way to do this would
be to include the edge (2, 6) in a new matching . This inclusion requires removal
of (1, 6), which can be compensated by inclusion of (1, 7) in the new matching.
This new matching Me= {(1, 7), (2, 6), (4, 8), (5, 9)) is shown in Figure 10.9c.
In general, we increase the size of a current matching M by constructing a
simple path from a free vertex in V to a free vertex in U whose edges are alternately
in E - M and in M. That is, the first edge of the path does not belong to M, the
second one does, and so on until the last edge that does not belong to M. Such a
path is called augmenting with respect to the matching M. For example, the path
2, 6, 1, 7 is an augmenting path with respect to the matching Mb in Figure 10.9b.
Since the length of an augmenting path is always odd, adding to the matching M
the path's edges in the odd-numbered positions and deleting from it the path's
edges in the even-numbered positions yields a matching with one more edge than
in M. Such a matching adjustment is called augmentation. Thus, in Figure 10.9,
the matching Mb was obtained by augmentation of the matching M0 along the
augmenting path 1, 6, and the matching Me was obtained by augmentation of the
matching Mb along the augmenting path 2, 6, 1, 7. Moving further, 3, 8, 4, 9, 5, 10
is an augmenting path for the matching Me (see Figure 10.9c). After adding to
Me the edges (3, 8), (4, 9), and (5, 10) and deleting (4, 8) and (5, 9), we obtain
the matching Md = { (1, 7), (2, 6), (3, 8), (4, 9), (5, 10)) shown in Figure 10.9d. The
r
10.3
I
Maximum Matching in Bipartite Graphs
v
u
7
(a)
Augmenting path: 1, 6
7
(b)
Augmenting path: 2, 6, 1, 7
(c)
Augmenting path: 3, 8, 4, 9, 5, 10
9
(d)
Maximum matching
FIGURE 10.9 Augmenting paths and matching augmentations
365
.......,
m
... '.· ~· ·I·.
I
'l.
!\ ' I;
366
Iterative Improvement
matching Md is not only a maximum matching but also peifect, i.e., a matching
that matches all the vertices of the graph.
Before we discuss an algorithm for finding an augmenting path, let us settle
the issue of what nonexistence of such a path means. According to the theorem
discovered by the French mathematician Claude Berge, it means the current
matching is maximal.
A matching M is maximal if and only if there exists no augmenting
path with respect toM.
THEOREM
PROOF If an augmenting path with respect to a matching M exists, then the size
of the matching can be increased by augmentation and hence cannot be maximal.
Let us prove the more difficult part: if no augmenting path with respect to a
matching M exists, then the matching is maximal. Assume that, on the contrary,
this is not the case for a certain matching M in a graph G. Let M' be a maximum
matching in G; by our assumption, the number of edges in M* is at least one
more than the number of edges in M, i.e., 1M* I> IMI. Consider the edges in the
symmetric difference M Ell M* = (M- M') U (M'- M), the set of all the edges
that are either in Morin M* but not in both. Note that IM*- Ml > IM- M*l
because IM'I > IMI by assumption. Let G' be the subgraph of G made up of
all the edges in M Ell M* and their endpoints. By definition of a matching, any
vertex in G' s; G can be incident to no more than one edge in M and no more
than one edge in M*. Hence each of the vertices in G' has degree 2 or less, and
therefore every connected component of G' is either a path or an even-length cycle
of alternating edges from M- M* and M'- M. Since IM'- Ml > IM- M*l and
the number of edges from M - M' and M* - M is the same for any even-length
cycle of alternating edges in G', there must exist at least one path of alternating
edges that starts and ends with an edge from M' - M. Hence, this is an augmenting
path for the matching M, which contradicts the assumption that no such path exists.
Ill
Our discussion of augmenting paths leads to the following general method
for constructing a maximum matching in a bipartite graph. Start with some initial
matching (e.g., the empty set). Find an augmenting path and augment the current
matching along this path. When no augmenting path can he found, terminate the
algorithm and return the last matching, which is maximal.
We now give a specific algorithm implementing this general template. We will
search for an augmenting path for a matching M by a BFS-like traversal of the
graph that starts simultaneously at all the free vertices in one of the sets V and U,
say V. (It would be logical to select the smaller of the two vertex sets, but we will
ignore this observation in the pseudocode below.) Recall that an augmenting path,
if it exists, is an odd-length path that connects a free vertex in V with a free vertex
in U and which, unless it consists of a single edge, "zigs" from a vertex in V to
another vertex' mate in U, then "zags" back to V along the uniquely defined edge
from M, and so on until a free vertex in U is reached. (Draw augmenting paths
10.3
Maximum Matching in Bipartite Graphs
367
for the matchings in Figure 10.9, for example.) Hence, any candidate to be such a
path must have its edges alternate in the pattern just described. This motivates the
following rules for labeling vertices during the BFS-like traversal of the graph.
Case 1 (the queue's front vertex w is in V) If u is a free vertex adjacent tow,
it is used as the other endpoint of an augmenting path; so the labeling stops and
augmentation of the matching commences. The augmenting path in question is
obtained by moving backwards along the vertex labels (see below) to alternately
add and delete its edges to and from the current matching. If u is not free and
connected to w by an edge not in M, label u with "w" unless it has been already
labeled.
Case 2 (the front vertex w is in U) In this case, w must be matched and we
label its mate in V with "w."
Here is a pseudocode of the algorithm in its entirety.
ALGORITHM
MaximumBipartiteMatching( G)
//Finds a maximum matching in a bipartite graph by a BFS-like traversal
//Input: A bipartite graph G = (V, U, E)
//Output: A maximum-cardinality matching Min the input graph
initialize set M of edges with some valid matching (e.g., the empty set)
initialize queue Q with all the free vertices in V (in any order)
while not Empty(Q) do
w <- Front(Q); Dequeue(Q)
ifw E V
for every vertex u adjacent to w do
if u is free
//augment
M <-MU(w, u)
V<-W
while vis labeled do
u <-vertex indicated by v's label; M <- M- (v, u)
v <- vertex indicated by u 's label; M <- M U (v, u)
remove all vertex labels
reinitialize Q with all free vertices in V
break //exit the for loop
else 1/u is matched
if (w, u) <;t M and u is unlabeled
label u with w
Enqueue(Q, u)
else 1/w E U (and matched)
label the mate v of w with "w"
Enqueue(Q, v)
return M //current matching is maximum
-.,
!
Iterative Improvement
368
2
2
7
7
Queue: 1 2 3
Queue: 1 2 3
1'
Augment from 6
8
6
2
2
7
7
Queue: 2 3
2
1
3
Queue:236814
1'1' 1' 1' 1'
Augment from 7
8
6
3
Queue: 3
3
4
Queue: 3 6 8 2 4 9
4
1' 1' 1' 1' 1'
Augment from 10
Queue: empty ~ maximum matching
FIGURE 10.10 Application of the maximum-cardinality-matching algorithm. The left
column shows a current matching and initialialized queue at the next
iteration's start; the right column shows the vertex labeling generated
by the algorithm before augmentation is performed. Matching edges are
shown in bold. Vertex labels indicate the vertices from which the labeling
is done. The discovered endpoint of an augmenting path is shaded and
labeled for clarity. Vertices deleted from the queue are indicated by the
t sign.
10.3
Maximum Matching in Bipartite Graphs
369
An application of this algorithm to the matching in Figure 10.9a is shown in
Figure 10.10. Note that the algorithm finds a maximum matching that differs from
the one in Figure 10.9.
How efficient is the maximum-matching algorithm? Each iteration except
the last one matches two previously free vertices-one from each of the sets V
and U. Therefore the total number of iterations cannot exceed [n/2J + 1, where
n = IV I + IU I is the number of vertices in the graph. The time spent on each
iteration is in O(n + m), where m =lEI is the number of edges in the graph. (This
assumes that the information about the status of each vertex-free or matched and
the vertex' mate if the latter-can be retrieved in constant time, e.g., by storing it in
an array.) Hence, the time efficiency of the algorithm is in O(n (n + m)). Hopcroft
and Karp [Hop73] showed how the efficiency can be improved to 0(.../fi(n + m))
by combining several iterations into a single stage to maximize the number of
edges added to the matching with one search.
We were concerned in this section with matching the largest possible number of vertex pairs in a bipartite graph. Some applications may require taking
into account the quality or cost of matching different pairs. For example, workers
may execute jobs with different efficiencies, or girls may have different preferences for their potential dance partners. It is natural to model such situations
by bipartite graphs with weights assigned to their edges. This leads to the problem of maximizing the sum of the weights on edges connecting matched pairs of
vertices. This problem is called maximum-weight matching. We encountered it
under a different name-the assignment problem-in Section 3.4. There are several sophisticated algorithms for this problem, which are much more efficient than
exhaustive search (see, e.g., [Pap82], [Gal86], [Ahu93]). We have to leave them
outside of our discussion, however, because of their complexity, especially for
general graphs.
-----Exercises 1 0 . 3 - - - - - - - - - - - - - - 1. For each matching shown below in bold, find an augmentation or explain why
no augmentation exists.
a.
b.
4
0
---,
'
'1.1•
I'!)
i I
I
"·····.·'·;.·.11•.
.~
' '
'
370
Iterative Improvement
2. Apply the maximum-matching algorithm to the following bipartite graph:
3
4
6
3. a. What is the largest and what is the smallest possible cardinality of a matching in a bipartite graph G = (V, U, E) with n vertices in each vertex set V
and U and at least n edges?
b. What is the largest and what is the smallest number of distinct solutions
the maximum-cardinality-matching problem can have for a bipartite graph
G = (V, U, E) with n vertices in each vertex set V and U and at least n
edges?
4. a. Hall's Marriage Theorem asserts that a bipartite graph G = (V, U, E) has a
matching that matches all vertices of the set V if and only if for each subset
S <; V, IR(S)I2': lSI where R(S) is the set of all vertices adjacent to a vertex
inS. Check this property for the following graph with (i) V = {1, 2, 3, 4}
and (ii) V = {5, 6, 7}.
4
b. You have to devise an algorithm that returns yes if there is a matching in
a bipartite graph G = ( V, U, E) that matches all vertices in V and returns
no otherwise. Would you base your algorithm on checking the condition
of Hall's Marriage Theorem?
5. Suppose there are five committees A, B, C, D, and E composedofsixpersons
a, b, c, d, e, and f as follows: committee A's members are band e; committee
B's members are b, d and e; committee C's members are a, c, d, e, and f;
committee D's members are b, d, and e; committee E's members are b and
e. Is there a system of distinct representatives, i.e., is it possible to select
a representative from each committee so that all the selected persons are
distinct?
6. Show how the maximum-cardinality-matching problem in a bipartite graph
can be reduced to the maximum-flow problem discussed in Section 10.2.
7. Consider the following greedy algorithm for finding a maximum matching in
a bipartite graph G = ( V, U, E). Sort all the vertices in nondecreasing order
10.4 The Stable Marriage Problem
371
of their degrees. Scan this sorted list to add to the current matching (initially
empty) the edge from the list's free vertex to an adjacent free vertex of the
lowest degree. If the list's vertex is matched or if there are no adjacent free
vertices for it, the vertex is simply skipped.
Does this algorithm always produce a maximum matching in a bipartite
graph?
8. Design a linear algorithm for finding a maximum matching in a tree.
9. Implement the maximum-matching algorithm of this section in the language
of your choice. Experiment with its performance on bipartite graphs with n
vertices in each of the vertex sets and randomly generated edges (in both
dense and sparse modes) to compare the observed running time with the
algorithm's theoretical efficiency.
10. Domino puzzle A domino is a 2-by-1 tile that can be oriented either horizontally or vertically. A tiling of a given board composed of 1-by-1 squares is
covering it with dominoes exactly without overlap. For each of the following
boards, determine if it can be tiled with dominoes. If you answer yes, indicate
a tiling; if you answer no, explain why.
a.
b.
c.
1 0.4 The Stable Marriage Problem
In this section we consider an interesting version of bipartite matching called the
stable marriage problem. Consider a set Y = {rn 1, rn 2, ... , rn,) of n men and a
set X= {wlo w2 , ... , wn) of n women. Each man has a preference list ordering
the women as potential marriage partners with no ties allowed. Similarly, each
woman has a preference list of the men, also with no ties. Examples of these two
sets of lists are given in Figure lO.lla and lO.llb. The same information can also he
presented by an n-by-n ranking matrix (see Figure 10.1lc). The rows and columns
of the matrix represent the men and women of the two sets, respectively. A cell
in row rn and column w contains two rankings: the first is the position (ranking)
of win them's preference list; the second is the position (ranking) of min thew's
l
372
Iterative Improvement
men's preferences
Bob:
Jim:
Tom:
2nd
Ann
Sue
Lea
1st
Lea
Lea
Sue
(a)
3rd
Sue
Ann
Ann
ranking matrix
women's preferences
Ann:
Lea:
Sue:
lst
2nd
Tom
Jim
Tom Bob
Jim
Tom
(b)
3rd
Bob
Jim
Bob
Ann
Bob
Jim
Torn
~
3,1
3,2
Lea
1,2
Sue
IT2l
2,1
2,1
3,3
~
(c)
FIGURE 10.11 Data for an instance of the stable marriage problem: (a) men's preference
lists; (b) women's preference lists; (c) ranking matrix (with the boxed cells
composing an unstable matching)
preference list. For example, the pair 3, 1 in Jim's row and Ann's column in the
matrix in Figure 10.11c indicates that Ann is Jim's tlmd choice while Jim is Ann's
first. (Which of these two ways to represent such information is better depends on
the task at hand. For example, it is easier to specify a match of the sets' elements
by using the ranking matrix, while the preference lists might be a more efficient
data structure for implementing a matching algorithm.)
A marriage matching M is a set of n (m, w) pairs whose members are selected
from disjoint n-element sets Y and X in a one-one fashion, i.e., each man m from Y
is paired with exactly one woman w from X and vice versa. (If we represent Y and X
as vertices of a complete bipartite graph with edges connecting possible marriage
partners, then a marriage matching is a perfect matching in such a graph.) A pair
(m, w), where mE Y, wE X, is said to be a blocking pair for a marriage matching
M if man m and woman w are not matched in M but they prefer each other to
their mates in M. For example, (Bob, Lea) is a blocking pair for the marriage
matching M = [(Bob,Ann), (Jim,Lea), (Tom,Sue)} (Figure lO.llc) because they
are not matched in M while Bob prefers Lea to Ann and Lea prefers Bob to Jim. A
marriage matching M is called stable if there is no blocking pair for it; otherwise,
Miscalled unstable. According to this definition, the marriage matching in Figure
lO.llc is unstable because Bob and Lea can drop their designated mates to join in
a union they both prefer. The stable marriage problem is to find a stable marriage
matching for men's and women's given preferences.
Surprisingly, this problem always has a solution. (Can you find it for the
instance in Figure 10.11 ?) It can be found by the following algorithm.
Stable marriage algorithm
Input: A set of n men and a set of n women along with rankings of the
women by each man and rankings of the men by each woman
Output: A stable marriage matching
Step 0 Start with all the men and women being free.
Step 1 While there are free men, arbitrarily select one of them and do the
following:
10.4 The Stable Marriage Problem
Free men:
Bob, Jim, Tom
Bob
Jim
Tom
Free men:
Jim, Tom
Bob
Jim
Tom
Freemen:
Jim, Tom
Bob
Jim
Tom
Freemen:
Tom
Free men:
Tom
Free men:
Bob
Bob
Jim
Tom
Bob
Jim
Tom
Bob
Jim
Tom
Ann
Lea
2,3
3, I
3, 2
[!d) 3,3
Sue
1.3
2, I
2, 1
1, 2
Ann
Lea
Sue
2,3
3, 1
3, 2
[!d) 3,3
hl
2, I
2, I
1, 2
Ann
Lea Sue
2,3
3, 1
3,2
[!d) 3,3
I, 3 ~
2, I
1, 2
Ann
Lea
Sue
2,3
3, I
3,2
[!d) 3,3
1, 3 ~
2, I
1, 2
Ann
Lea
Sue
2,3
3, 1
3, 2
1,2 3,3
I, 3 ~
~ 1, 2
Ann
Lea
~
1, 2 3,3
1,3 ~
~ I, 2
3, I
3,2
373
Bob proposed to Lea
Lea accepted
Jim proposed to Lea
Lea rejected
Jim proposed to Sue
Sue accepted
Tom proposed to Sue
Sue rejected
Tom proposed to Lea
Lea replaced Bob with Torn
Sue
Bob proposed to Ann
Ann accepted
FIGURE 10.12 Application of the stable marriage algorithm. An accepted proposal is
indicated by a boxed cell; a rejected proposal is shown by an underlined
cell.
Proposal The selected free man m proposes to w, the next woman
on his preference list (who is the highest-ranked woman who has not
rejected him before).
Response If w is free, she accepts the proposal to be matched with
m. If she is not free, she compares m with her current mate. If she
prefers m to him, she accepts m's proposal, making her former mate
free; otherwise, she simply rejects m 's proposal, leaving m free.
Step 2 Return the set of n matched pairs.
Before we analyze this algorithm, it is useful to trace it on some input. Such
an example is presented in Figure 10.12.
Let us discuss properties of the stable marriage algorithm.
374
Iterative Improvement
THEOREM The stable marriage algorithm terminates after no more than n 2
iterations with a stable marriage output.
PROOF The algorithm starts with n men having the total of n 2 women on their
ranking lists. On each iteration, one man makes a proposal to a woman. This
reduces the total number of women to whom the men can still propose in the
future because no man proposes to the same woman more than once. Hence, the
algorithm must stop after no more than n 2 iterations.
Let us now prove that the final matching M is a stable marriage matching.
Since the algorithm stops after all then men are one-one matched to then women,
the only thing that needs to be proved is the stability of M. Suppose, on the
contrary, that M is unstable. Then there exists a blocking pair of a man m and a
woman w who are unmatched in M and such that both m and w prefer each other
to the persons they are matched with in M. Since m proposes to every woman on
his ranking list in decreasing order of preference and w precedes m 's match in M,
m must have proposed to w on some iteration. Whether w refused m 's proposal or
accepted it but replaced him on a subsequent iteration with a higher-ranked match,
w's mate in M must be higher on w's preference list than m because the rankings
of the men matched to a given woman may only improve on each iteration of the
algorithm. This contradicts the assumption that w prefers m to her final match
inM.
•
The stable marriage algorithm has a notable shortcoming. It is not "gender
neutral." In the form presented above, it favors men's preferences over women's
preferences. We can easily see it by tracing the algorithm on the following instance
of the problem:
man 1
man2
woman 1
1,2
2,1
woman 2
2,1
1,2
The algorithm obviously yields the stable matching M ~{(man 1, woman 1), (man
2, woman 2)). In this matching, both men are matched to their first choices, which
is not the case for the women. One can prove that the algorithm always yields a
stable matching that is man-optimal: it assigns to each man the highest-ranked
woman possible under any stable marriage. Of course, this gender bias can be
reversed, but not eliminated, by reversing the roles played by men and women
in the algorithm, i.e., by making women propose and men accept or reject their
proposals.
There is another important corollary to the fact that the stable marriage
algorithm always yields a gender-optimal stable matching. It is easy to prove
that a man (woman)-optimal matching is unique for a given set of participant
preferences. Therefore the algorithm's output does not depend on the oroer in
which the free men (women) make their proposals. Consequently, we can use any
10.4 The Stable Marriage Problem
375
data structure we might prefer-e.g., a queue or a stack-for representing this set
with no impact on the algorithm's outcome.
The notion of the stable matching as well as the algorithm discussed above was
introduced by D. Gale and L. S. Shapley in the paper titled "College Admissions
and the Stability of Marriage" [Gal62]. I do not know which of the two applications
mentioned in the title you would consider more important. The point is that
stability is a matching property that can be desirable in a variety of applications.
For example, it has been used for many years in the United States for matching
medical-school graduates with hospitals for residency training. For a brief history
of this application and an in-depth discussion of the stable marriage problem and
its extensions, see the monograph by Gusfield and Irwing [Gus89].
-----Exercises 1 0 . 4 - - - - - - - - - - - - - - 1. Consider an instance of the stable marriage problem given by the ranking
matrix
rx
f3
y
A
1, 3
3, 1
2,2
c
B
2,2 3, 1
1, 3 2, 2
3, 1 1, 3
For each of its marriage matchings, indicate whether it is stable or not. For the
unstable matchings, specify a blocking pair. For the stable matchings, indicate
whether they are man-optimal, woman-optimal, or neither. (Assume that the
Greek and Roman letters denote the men and women, respectively.)
2. Design a simple algorithm for checking whether a given marriage matching is
stable and determine its time efficiency class.
3. Find a stable marriage matching for the instance given in Problem 1 by applying the stable marriage algorithm
a. in its men-proposing version.
b. in its women-proposing version.
4. Find a stable marriage matching for the instance defined by the following
ranking matrix:
rx
A
B
1, 3 2, 3
1, 4 4, 1
2,2 1, 4
4, 1 2, 2
c
D
3, 2 4, 3
f3
3,4 2, 2
y
3,3 4, 1
8
3, 1 1, 4
5. Determine the time-efficiency class of the stable marriage algorithm
a. in the worst case.
b. in the best case.
I
fr
rl
376
Iterative Improvement
6. Prove that a man-optimal stable marriage set is always unique. Is it also true
for a woman-optimal stable marriage matching?
7. Prove that in the man-optimal stable matching, each woman has the worst
partner that she can have in any stable marriage matching.
8. Implement the stable marriage algorithm given in Section 10.4 so that its
running time is in O(n 2 ). Run an experiment to ascertain its average-case
efficiency.
9. Write a report on the college admission problem (residents-hospitals assignment) that generalizes the stable marriage problem.
10. Consider the problem of the roommates, which is related to but more difficult
than the stable marriage problem: "An even number of boys wish to divide up
into pairs of roommates. A set of pairings is called stable if under it there are
no two boys who are not roommates and who prefer each other to their actual
roommates." [Gal62] Give an instance of this problem that does not have a
stable pairing.
SUMMARY
Bl!l
The iterative improvement technique involves finding a solution to an optimiza-
tion problem by generating a sequence of feasible solutions with improving
values of the problem's objective function. Each subsequent solution in such
a sequence typically involves a small, localized change in the previous feasible
solution. When no such change improves the value of the objective function,
the algorithm returns the last feasible solution as optimal and stops.
'" Important problems that can be solved exactly by iterative improvement
algorithms include linear programming, maximizing the flow in a network,
and matching the maximum possible number of vertices in a graph.
11
The simplex method is the classic method for solving the general linear programming problem. It works by generating a sequence of adjacent extreme
points of the problem's feasible region with improving values of the objective
function.
Iii
The maximum-flow problem asks to find the maximal flow possible in a network, a weighted directed graph with a source and a sink.
'" The Ford-Fulkerson method is a classic template for solving the maximumflow problem by the iterative-improvement approach. The shortestaugmenting-path method implements this idea by labeling network vertices
in the breadth-first search manner.
'" The Ford-Fulkerson method also finds a minimum cut in a given network.
I
Summary
II
Ill
377
A maximum cardinality matching is the largest subset of edges in a graph such
that no two edges share the same vertex. For a bipartite graph, it can be found
by a sequence of augmentations of previously obtained matchings.
The stable marriage problem is to find a stable matching for elements of two nelement sets based on given matching preferences. This problem always has
a solution that can be found by the stable marriage algorithm.
lr
It
11
Limitations of
Algorithm Power
Intellect distinguishes between the possible and the impossible; reason
distinguishes between the sensible and the senseless. Even the possible
can be senseless.
-Max Born (1882-19701, My Life and My Views, 1968
In the preceding chapters of this book, we encountered dozens of algorithms
I for solving a variety of different problems. A fair assessment of algorithms as
problem-solving tools is inescapable: they are very powerful instruments, especially when they are executed by modern computers. But the power of algorithms
is not unlimited, and its limits are the subject of this chapter. As we shall see, some
problems cannot be solved by any algorithm. Other problems can be solved algorithmically but not in polynomial time. And even when a problem can be solved
in polynomial time by some algorithms, there are usually lower bounds on their
efficiency.
We start, in Section 11.1, with methods for obtaining lower bounds, i.e., estimates on a minimum amount of work needed to solve a problem. In general,
obtaining a nontrivial lower bound even for a simple-sounding problem is a very
difficult task. As opposed to ascertaining the efficiency of a particular algorithm,
the task here is to establish a limit on the efficiency of any algorithm, known or
unknown. This also necessitates a careful description of the operations such algorithms are allowed to perform. If we fail to define carefully the "rules of the game,"
so to speak, our claims may end up in the large dustbin of impossibility-related
statements as, for example, the one made by the celebrated British physicist Lord
Kelvin in 1895: "Heavier-than-air flying machines are impossible."
Section 11.2 discusses decision trees. This technique allows us, among other
applications, to establish lower bounds on the efficiency of comparison-based
algorithms for sorting and for searching in sorted arrays. As a result, we will be
able to answer such questions as whether it is possible to invent a faster sorting
algorithm than mergesort and whether binary search is the fastest algorithm for
379
l
Limitations of Algorithm Power
380
searching in a sorted array. (What does your intuition tell you the answers to
these questions will turn out to be?) Incidentally, decision trees are also a great
vehicle for directing us to a solution of some puzzles, such as the fake-coin problem
discussed in Section 5.5.
Section 11.3 deals with the question of intractability: which problems can
and cannot be solved in polynomial time. Thls well-developed area of theoretical
computer science is called computational complexity theory. We present the basic
elements of this theory and discuss informally such fundamental notions as P, NP,
and NP-complete problems, including the most important unresolved question of
theoretical computer science about the relationship between P and NP problems.
The last section of this chapter deals with numerical analysis. Thls branch
of computer science concerns algorithms for solving problems of "continuous"
mathematics-solving equations and systems of equations, evaluating such functions as sin x and In x, computing integrals, and so on. The nature of such problems
imposes two types of limitations. First, most cannot be solved exactly. Second,
solving them even approximately requires dealing with numbers that can be represented in a digital computer with only a limited level of precision. Manipulating
approximate numbers without proper care can lead to very inaccurate results. We
will see that even solving a basic quadratic equation on a computer poses significant difficulties that require a modification of the canonical formula for the
equation's roots.
11.1
Lower-Bound Arguments
We can look at the efficiency of an algorithm two ways. We can establish its asymptotic efficiency class (say, for the worst case) and see where this class stands with
respect to the hierarchy of efficiency classes outlined in Section 2.2. For example, selection sort, whose efficiency is quadratic, is a reasonably fast algorithm,
whereas the algorithm for the Tower of Hanoi problem is very slow because its efficiency is exponential. One can argue, however, that this comparison is akin to the
proverbial comparison of apples to oranges because these two algorithms solve
different problems. The alternative and possibly "fairer" approach is to ask how
efficient a particular algorithm is with respect to other algorithms for the same
problem. Seen in this light, selection sort has to be considered slow because there
are 0 (n log n) sorting algorithms; the Tower of Hanoi algorithm, on the other
hand, turns out to be the fastest possible for the problem it solves.
When we want to ascertain the efficiency of an algorithm with respect to other
algorithms for the same problem, it is desirable to know the best possible efficiency
any algorithm solving the problem may have. Knowing such a lower bound can
tell us how much improvement we can hope to achieve in our quest for a better
algorithm for the problem in question. If such a bound is tight, i.e., we already
know an algorithm in the same efficiency class as the lower bound, we can hope
for a constant-factor improvement at best. If there is a gap between the efficiency
I.
II
I
11.1
Lower-Bound Arguments
381
of the fastest algorithm and the best lower bound known. the door for possible
improvement remains open: either a faster algorithm matching the lower bound
could exist or a better lower bound could be proved.
In this section. we present several methods for establishing lower bounds and
illustrate them with specific examples. As we did in analyzing the efficiency of
specific algorithms in the preceding chapters, we should distinguish between a
lower-bound class and a minimum number of times a particular operation needs
to be executed. As a rule, the second problem is more difficult than the first.
For example, we can immediately conclude that any algorithm for finding the
median of n numbers must be in !J(n) (why?), but it is not simple at all to prove
that any comparison-based algorithm for this problem must do at least 3(n- 1)/2
comparisons in the worst case (for odd n ).
Trivial lower Bounds
The simplest method of obtaining a lower-bound class is based on counting the
number of items in the problem's input that must be processed and the number of
output items that need to be produced. Since any algorithm must at least "read"
all the items it needs to process and "write" all its outputs, such a count yields a
trivial lower bound. For example, any algorithm for generating all permutations
of n distinct items must be in Q(n!) because the size of the output is n!. And this
bound is tight because good algorithms for generating permutations spend a
constant time on each of them except the initial one (see Section 5.4).
As another example, consider the problem of evaluating a polynomial of
degree n
at a given point x, given its coefficients an, a11 _ 1, ... , a0 . It is easy to see that all the
coefficients have to be processed by any polynomial-evaluation algorithm. Indeed,
if it were not the case, we could change the value of an unprocessed coefficient,
which would change the value of the polynomial at a nonzero point x. This means
that any such algorithm must be in Q (n). This lower bound is tight because both
the right-to-left evaluation algorithm (Problem 2 in Exercises 6.5) and Horner's
rule (Sectiou 6.5) are both linear.
In a similar vein, a trivial lower bound for computing the product of two nby-n matrices is Q (n 2) because any such algorithm has to process 2n 2 elements in
the input matrices and generate n 2 elements of the product. It is still unknown,
however, whether this bound is tight.
Trivial lower bounds are often too low to be useful. For example, the trivial
bound for the traveling salesman problem is Q (n2 ), because its input is n (n - l) /2
intercity distances and its output is a list of n + 1 cities making up an optimal tour.
But this bound is all but useless since there is no known algorithm for this problem
with the running time being a polynomial function of any degree.
~II
cl -~- !!
if 1 ;
·
j ·.•·.·.· .·.·
•
'·· '··
•.•.
.•..
~-
:
II
I
382
Limitations of Algorithm Power
There is another obstacle to deriving a meaningful lower bound by this
method. It lies in determining which part of an input must be processed by any
algorithm solving the problem in question. For example, searching for an element of a given value in a sorted array does not require processing all its elements
(why?). As another example, consider the problem of determining connectivity of
an undirected graph defined by its adjacency matrix. It is plausible to expect that
any such algorithm would have to check the existence of each of the n(n- 1)/2
potential edges, but the proof of this fact is not trivial.
Information-Theoretic Arguments
While the approach outlined above takes into account the size of a problem's
output, the information-theoretical approach seeks to establish a lower bound
based on the amount of information it has to produce. Consider, as an example,
the well-known game of deducing a positive integer between 1 and n selected
by somebody by asking that person questions with yes/no answers. The amount of
uncertainty that any algorithm solving this problem has to resolve can be measured
by rlog2 n l, the number of bits needed to specify a particular number among the
n possibilities. We can think of each question (or, to be more accurate, an answer
to each question) as yielding at most one bit of information about the algorithm's
output, i.e., the selected number. Consequently, any such algorithm will need at
least rlog2 n l such steps before it can determine its output in the worst case.
The approach we just exploited is called the information-theoretic argument
because of its connection to information theory. It has proved to be quite useful
for finding the so-called information-theoretic lower bounds for many problems
involving comparisons, including sorting and searching. Its underlying idea can be
realized much more precisely through the mechanism of decision trees. Because
of the importance of this technique, we discuss it separately and in more detail in
Section 11.2.
Adversary Arguments
Let us revisit the same game of "guessing" a number used to introduce the idea of
an information-theoretic argument. We can prove that any algorithm that solves
this problem must ask at least rlog2 n l questions in its worst case by playing the
role of a hostile adversary who wants to make an algorithm ask as many questions
as possible. The adversary starts by considering each of the numbers between
1 and n as being potentially selected. (This is cheating, of course, as far as the
game is concerned, but not as a way to prove our assertion.) After each question,
the adversary gives an answer that leaves him with the largest set of numbers
consistent with this and all the previously given answers. (This strategy leaves
him with at least one half of the numbers he had before his last answer.) If an
algorithm stops before the size of the set is reduced to one, the adversary can
exhibit a number that could be a legitimate input the algorithm failed to identify.
li·,
11.1
Lower-Bound Arguments
383
It is a simple technical matter now to show that one needs flog2 n1 iterations to
shrink an n-element set to a one-element set by halving and rounding up the size
of the remaining set. Hence, at least flog 2 n1 questions need to be asked by any
algorithm in the worst case.
This example illustrates the adversary method for establishing lower bounds.
It is based on following the logic of a malevolent but honest adversary: the malevolence makes him push the algorithm down the most time-consuming path, while
his honesty forces him to stay consistent with the choices already made. A lower
bound is then obtained by measuring the amount of work needed to shrink a set
of potential inputs to a single input along the most time-consuming path.
As another example, consider the problem of merging two sorted lists of size n
a1 < a2 < ... < an and b1 < b2 < ... < b11
into a single sorted list of size 2n. For simplicity, we assume that all the a's and
b's are distinct, which gives the problem a unique solution. We encountered this
problem when discussing mergesort in Section 4.1. Recall that we did merging by
repeatedly comparing the first elements in the remaining lists and outputting the
smaller among them. The number of key comparisons in the worst case for this
algorithm for merging is 2n - 1.
Is there an algorithm that can do merging faster? The answer turns out to be
no. Knuth ([Knuiii], p. 198) quotes the following adversary method for proving
that 2n - 1 is a lower bound on the number of key comparisons made by any
comparison-based algorithm for this problem. The adversary will employ the
following rule: reply true to the comparison a; < b1 if and only if i < j. This will
force any correct merging algorithm to produce the only combined list consistent
with this rule:
To produce this combined list, any correct algorithm will have to explicitly compare 2n - 1 adjacent pairs of its elements, i.e., b1 to
a 1 to b2 , and so on. If one
of these comparisons has not been made, e.g., has not been compared to b2 , we
can transpose these keys to get
a,
a,
which is consistent with all the comparisons made but cannot be distinguished
from the correct configuration given above. Hence, 2n - 1 is indeed a lower bound
for the number of key comparison needed for any merging algorithm.
Problem Reduction
We have already encountered the problem-reduction approach in Section 6.6.
There, we discussed getting an algorithm for problem P by reducing it to another
problem Q solvable with a known algorithm. A similar reduction idea can be used
for finding a lower bound. To show that problem P is at least as hard as another
.......,
I
384
Limitations of Algorithm Power
TABLE 11.1 Problems often used for establishing lower bounds
by problem reduction
Problem
Lower bound
Tightness
sorting
iJ(nlogn)
searching in a sorted array
iJ(Iog n)
element uniqueness problem
iJ(n logn)
yes
yes
yes
multiplication of n-digit integers
iJ(n)
multiplication of square matrices
iJ (n')
unknown
unknown
problem Q with a known lower bound, we need to reduce Q to P (not P to Q!).
In other words, we should show that an arbitrary instance of problem Q can be
transformed (in a reasonably efficient fashion) to an instance of problem P, so
any algorithm solving P would solve Q as well. Then a lower bound for Q will be
a lower bound for P. Table 11.1lists several important problems that are often
used for this purpose.
We will establish the lower bounds for sorting and searching in the next section. The element uniqueness problem asks whether there are duplicates among n
given numbers. (We encountered this problem in Sections 2.3 and 6.1.) The proof
of the lower bound for this seemingly simple problem is based on a very sophisticated mathematical analysis that is well beyond the scope of this book (see, e.g.,
[PreSS] for a rather elementary exposition). As to the last two algebraic problems
in Table 11.1, there the lower bounds quoted are trivial, but whether they can be
improved remains unknown.
As an example of establishing a lower bound by reduction, let us consider
the Euclidean minimum spanning tree problem: given n points in the Cartesian
plane, construct a tree of minimum total length whose vertices are the given
points. As a problem with a known lower bound, we use the element uniqueness
problem. We can transform any set x 1 , x 2 , ... , xn of n real numbers into a set of
n points in the Cartesian plane by simply adding 0 as the points' y coordinate:
(x 1, 0), (x2 , 0), ... , (x,,O). LetT be a minimum spanning tree found for this set of
points. Since T must contain a shortest edge, checking whether T contains a zerolength edge will answer the question about uniqueness of the given numbers. This
reduction implies that Q (n log n) is a lower bound for the Euclidean minimum
spanning tree problem, too.
Since the final results about the complexity of many problems are not known,
the reduction technique is often used to compare the relative complexity of problems. For example, the formulas
X ·y=
(x + y ) 2 - (x - y )2
4
and
x 2 =x · x
11.1
Lower-Bound Arguments
385
show that the problems of computing the product of two n-digit integers and
squaring ann-digit integer belong to the same complexity class, despite the latter
being seemingly simpler than the former.
There are several similar results for matrix operations. For example, multiplying two symmetric matrices turns out to be in the same complexity class as
multiplying two arbitrary square matrices. This result is based on the observation
that not only the former problem is a special case of the latter one, but we also
can reduce the problem of multiplying two arbitrary square matrices of order n,
say A and B, to the problem of multiplying two symmetric matrices
~]andY=[~ BT]
0
,
where AT and BT are the transpose matrices of A and B (i.e., AT[i, j] = A[j, i] and
BT[i, j] = B[.i, i]), respectively, and 0 stands for the n-by-n matrix whose elements
are all zeros. Indeed,
XY=
0 J,
0
B BoT]=[AoB ATBT
[AT0 A][O
from which the needed product AB can be easily extracted. (True, we will have
to multiply matrices twice the original size, but this is just a minor technical
complication with no impact on the complexity classes.)
Though such results are interesting, we will encounter even more important
applications of the reduction approach to comparing problem complexity in Section 11.3.
-----Exercises 11.1 - - - - - - - - - - - - - - 1. Prove that any algorithm solving the alternating-disk puzzle (Problem 11 in
Exercises 3.1) must make at least n(n + 1)/2 moves to solve it. Is this lower
bound tight?
2. Prove that the classic recursive algorithm for the Tower of Hanoi puzzle
(Section 2.4) makes the minimum number of disk moves needed to solve it.
3. Find a trivial lower-bound class for each of the following problems and indicate, if you can, whether this bound is tight.
a. finding the largest element in an array
b. checking completeness of a graph represented by its adjacency matrix
c. generating all the subsets of an-element set
d. determining whether n given real numbers are all distinct
4. Consider the problem of identifying a lighter fake coin among n identicallooking coins with the help of a balance scale. Can we use the same information-theoretic argument as the one in the text for the number of questions in
l
386
Limitations of Algorithm Power
the guessing game to conclude that any algorithm for identifying the fake will
need at least rlogz nl weighings in the worst case?
5. Prove that any comparison-based algorithm for finding the largest among n
given numbers must make n - 1 comparisons in the worst case.
6. Find a tight lower bound for sorting an array by exchanging its adjacent
elements.
7. Give an adversary-argument proof that the time efficiency of any algorithm
that checks connectivity of a graph with n vertices is in Q(n 2 ), provided the
only operation allowed for an algorithm is to inquire about the presence of
an edge between two vertices of the graph. Is this lower bound tight?
8. What is the minimum number of comparisons needed for a comparison-based
sorting algorithm to merge any two sorted lists of sizes n and n + 1 elements,
respectively? Prove the validity of your answer.
9. Find the product of matrices A and B through a transformation to a product
of two symmetric matrices if
A=[;
-;1 JandB=[~1 ;J.
10. a. Can we use this section's formulas that indicate the complexity equivalence of multiplication and squaring of integers to show the complexity
equivalence of multiplication and squaring of square matrices?
b. Show that multiplication of two matrices of order n can be reduced to
squaring a matrix of order 2n.
11. Find a tight lower-bound class for the problem of finding two closest numbers
among n real numbers xi> x2 , ... , x/1'
11.2 Decision Trees
Many important algorithms, especially those for sorting and searching, work by
comparing items of their inputs. We can study the performance of such algorithms
with a device called the decision tree. As an example, Figure 11.1 presents a
decision tree of an algorithm for finding a minimum of three numbers. Each
internal node of a binary decision tree represents a key comparison indicated
in the node, e.g., k < k'. The node's left subtree contains the information about
subsequent comparisons made if k < k 1, while its right subtree does the same for
the case of k > k'. (For the sake of simplicity, we assume throughout this section
that all input items are distinct.) Each leaf represents a possible outcome of the
algorithm's run on some input of size n. Note that the number of leaves can be
greater than the number of outcomes because, for some algorithms, the same
outcome can be arrived at through a different chain of comparisons. (Tbis happens
11.2
Decision Trees
387
a<b
no
FIGURE 11.1 Decision tree for finding a minimum of three numbers
to be the case for the decision tree in Figure 11.1.) An important point is that the
number of leaves must be at least as large as the number of possible outcomes.
The algorithm's work on a particular input of size n can be traced by a path from
the root to a leaf in its decision tree, and the number of comparisons made by the
algorithm on such a run is equal to the number of edges in this path. Hence, the
number of comparisons in the worst case is equal to the height of the algorithm's
decision tree.
The central idea behind this model lies in the observation that a tree with a
given number of leaves, which is dictated by the number of possible outcomes, has
to be tall enough to have that many leaves. Specifically, it is not difficult to prove
that for any binary tree with /leaves and height h,
(11.1)
Indeed, a binary tree of height h with the largest number of leaves has all its leaves
on the last level (why?). Hence, the largest number of leaves in such a tree is 2".
In other words, 2":::: I, which immediately implies (11.1).
Inequality (11.1) puts a lower bound on the heights of binary decision trees
and hence the worst-case number of comparisons made by any comparison-based
algorithm for the problem in question. Such a bound is called the informationtheoretic lower bound (see Section 11.1). We illustrate this technique below on
two important problems: sorting and searching in a sorted array.
Decision Trees for Sorting Algorithms
Most sorting algorithms are comparison-based, i.e., they work by comparing elements in a list to be sorted. Moreover, with the notable exception of binary
insertion sort (Problem 9 in Exercises 5.1), comparing two elements is the basic
operation of such algorithms. Therefore, by studying properties of decision trees
m
. · .l)t.~ .-·
n
l
•.
Limitations of Algorithm Power
388
r
abc
yes
0
no
abc
abc
yes
a<C
abc
yes
b<c
no
yes
no
b<C
cba
bac
cba
yes
no
a<c
no
yes
b<a
FIGURE 11.2 Decision tree for the three-element selection sort. A triple above a node
indicates the state of the array being sorted. Note the two redundant
comparisons b <a with a single possible outcome because of the results
of some previously made comparisons.
for comparison-based sorting algorithms, we can derive important lower bounds
on time efficiencies of such algorithms.
We can interpret an outcome of a sorting algorithm as finding a permutation
of the element indices of an input list that puts the list's elements in ascending
order. For example, for the outcome a < c < b obtained by sorting a list a, b, c
(see Figure 11.2), the permutation in question is 1, 3, 2. Hence, the number of
possible outcomes for sorting an arbitrary n-element list is equal ton!.
Inequality (11.1) implies that the height of a binary decision tree for any
comparison-based sorting algorithm and hence the worst -case number of comparisons made by such an algorithm cannot be less than [log2 11 !l:
Cwm·.,,(n) 2': rlog 2 n!l
(11,2)
Using Stirling's formula for n !, we get
;;,-"
log 11 log2 2n
r1og 2 n !l "' log2 -v ~.-n 11 (11 1e l = 11 log2 11 - n log2 e + - 2- +
"'n log2 n.
2
2
In other words, about n log2 n comparisons are necessary to sort an arbitrary
n-element list by any comparison-based sorting algorithm. Note that mergesort
makes about this number of comparisons in its worst case and hence is asymptotically optimal. This also implies that the asymptotic lower bound n log2 n is tight
and therefore cannot be substantially improved. We should point out, however,
11.2
Decision Trees
389
abc
yes
I
<2>
no
I
abc
yes
bac
no
b<c
yes
acb
yes
B<C
bca
no
yes
b<c
no
FIGURE 11.3 Decision tree for the three-element insertion sort
that the lower bound of rlog 2 nil can be improved for some values of n. For example, [log2 1211 = 29, but it has been proved that 30 comparisons are necessary
(and sufficient) to sort an array of 12 elements in the worst case.
We can also use decision trees for analyzing the average-case behavior of
a comparison"based sorting algorithm. We can compute the average number of
comparisons for a particular algorithm as the average depth of its decision tree's
leaves, i.e., as the average path length from the root to the leaves. For example,
for the three-element insertion sort whose decision tree is given in Figure 11.3,
this number is (2 + 3 + 3 + 2 + 3 + 3)/6 = 2~.
Under the standard assumption that all n! outcomes of sorting are equally
likely, the following lower bound on the average number of comparisons Ca,g
made by any comparison-based algorithm in sorting an n-element list has been
proved:
(11.3)
As we saw earlier, this lower bound is about n log2 n . You might be surprised
that the lower bounds for the average and worst cases are almost identical. Remember, however, that these bounds are obtained by maximizing the number of
comparisons made in the average and worst cases, respectively. For individual
sorting algorithms, the average-case efficiencies can, of course, be significantly
better than their worst-case efficiencies.
390
Limitations of Algorithm Power
Decision Trees for Searching a Sorted Array
In this section, we shall see how decision trees can be used for establishing lower
bounds on the number of key comparisons in searching a sorted array of n keys:
A[O] < A[l] < ... < A[n -1]. The principal algorithm for this problem is binary
search. As we saw in Section 4.3, the number of comparisons made by binary
search in the worst case, c~sorst (n), is given by the formula
(11.4)
We will use decision trees to determine whether this is the smallest possible
number of comparisons.
Since we are dealing here with three-way comparisons in which search key K is
compared with some element A[i] to see whether K < A[i], K = A[i], or K > A[i],
it is natural to try using ternary decision trees. Figure 11.4 presents such a tree for
the case of n = 4. The internal nodes of that tree indicate the array's elements being
compared with the search key. The leaves indicate either a matching element in
the case of a successful search or a found interval that the search key belongs to
in the case of an unsuccessful search.
We can represent any algorithm for searching a sorted array by three-way
comparisons with a ternary decision tree similar to that of Figure 11.4. For an
array of n elements, all such decision trees will have 2n +!leaves (n for successful
searches and n + 1 for unsuccessful ones). Since the minimum height h of a ternary
tree with l leaves is flog 3 1l, we get the following lower bound on the number of
worst-case comparisons:
>
<
.-----C A I01
<
>
<
FIGURE 11.4 Ternary decision tree for binary search in a four-element array
11.2
391
Decision Trees
This lower bound is smaller than riog 2 (n + 1)1, the number of worst-case
comparisons for binary search, at least for large values of n (and smaller than or
equal to [log2 (n + 1)1 for every positive integer n-see Problem 7 in the exercises).
Can we prove a better lower bound, or is binary search far from being optimal?
The answer turns out to be the former. To obtain a better lower bound, we should
consider binary rather than ternary decision trees such as the one in Figure 11.5.
Internal nodes in such a tree correspond to the same three-way comparisons
as before, but they also serve as terminal nodes for successful searches. Leaves
therefore represent only unsuccessful searches, and there are n + 1 of them for
searching ann-element array.
As comparison of the decision trees of Figures 11.4 and 11.5 illustrates, the
binary decision tree is simply the ternary decision tree with all the middle subtrees
eliminated. Applying inequality (11.1) to such binary decision trees immediately
yields
(11.5)
This inequality closes the gap between the lower bound and the number of worstcase comparisons made by binary search, which is also [log2 (n + 1)l A much
more sophisticated analysis (see, e.g., [Knulll], Section 6.2.1) shows that under the
standard assumptions about searches, binary search makes the smallest number
of comparisons on the average, as well. The average number of comparisons made
by this algorithm turns out to be about log2 n- 1 and log2 (n + 1) for successful
and unsuccessful searches, respectively.
All!
<
AIOI
<
Al21
>
<
Al31
<
FIGURE 11.5 Binary decision tree for binary search in a four-element array
392
Limitations of Algorithm Power
-----Exercises 11
1. Prove by mathematical induction that
a. h :> rlog 2 /l for any binary tree with height hand the number of leaves I.
b. h :> [log3 /l for any ternary tree with height hand the number of leaves I.
2. Consider the problem of finding the median of a three-element set fa. b, c).
a. What is the information-theoretic lower bound for comparison-based algorithms solving this problem?
b. Draw a decision tree for an algorithm solving this problem.
c. If the worst-case number of comparisons in your algorithm is greater
than the information-theoretic lower bound, do you think an algorithm
matching the lower bound exists? (Either find such an algorithm or prove
its impossibility.)
3. Draw a decision tree and find the number of key comparisons in the worst
and average cases for
a. the three-element basic bubble sort.
b. the three-element enhanced bubble sort (which stops if no swaps have been
made on its last pass).
4. Design a comparison-based algorithm for sorting a four-element array with
the smallest number of element comparisons possible.
5. Design a comparison-based algorithm for sorting a five-element array with
seven comparisons in the worst case.
6. Draw a binary decision tree for searching a four-element sorted list by sequential search.
7. Compare the two lower bounds for searching a sorted array-rlog3(2n + 1)1
and rlog2(n + 1)1-to show that
a. [log3(2n + 1)1 :s rlog2(n + 1)1 for every positive integer n.
b. rlog3(2n + 1)1 < pog2(n + 1)1 for every positive integer n )>no.
8. Advanced fake-coin problem There are n :> 3 coins identical in appearance;
either all are genuine or exactly one of them is fake. It is unknown whether
the fake coin is lighter or heavier than the genuine one. You have a balance
scale with which you can compare any two sets of coins. That is, by tipping to
the left, to the right, or staying even, the balance scale will tell whether the
sets weigh the same or which of the sets is heavier than the other, but not by
how much. The problem is to find whether all the coins are genuine and, if
not, to find the fake coin and establish whether it is lighter or heavier than the
genuine ones.
a. Prove that any algorithm for this problem must make at least rlog3 (2n + 1)1
weighings in the worst case.
11.3
P. NP, and NP·complete Problems
393
b. Draw a decision tree for an algorithm that solves this problem for n = 3
coins in two weighings.
c. Prove that there exists no algorithm that solves the problem for 11 = 4 coins
in two weighings.
d. Draw a decision tree for an algorithm that solves the problem for 11 = 4
coins in two weighings by using an extra coin known to be genuine.
e. Draw a decision tree for an algorithm that solves the classic version of
the problem-that for n = 12 coins in three weighings (with no extra coins
being used).
9. A tournament tree is a complete binary tree reflecting results of a "knockout
tournament": its leaves represent n players entering the tournament, and
each internal node represents a winner of a match played by the players
represented by the node's children. Hence, the winner of the tournament is
represented by the root of the tree.
a. What is the total number of games played in such a tournament?
b. How many rounds are there in such a tournament?
c. Design an efficient algorithm to determine the second-best player using
the information produced by the tournament. How many extra games does
your algorithm require?
10. Jigsaw puzzle A jigsaw puzzle contains n pieces. A "section" of the puzzle is
a set of one or more pieces that have been connected to each other. A "move"
consists of connecting two sections. What algorithm will minimize the number
of moves required to complete the puzzle?
11.3 P, IVP, and IVP-complete Problems
In the study of the computational complexity of problems, the first concern of both
computer scientists and computing professionals is whether a given problem can
be solved in polynomial time by some algorithm.
DEFINITION 1 We say that an algorithm solves a problem in polynomial time
if its worst-case time efficiency belongs to O(p(n)) where p(n) is a polynomial
of the problem's input size n. (Note that since we are using big·oh notation
here, problems solvable in, say, logarithmic time are solvable in polynomial time
as well.) Problems that can be solved in polynomial time are called tractable,
problems that cannot be solved in polynomial time are called intractable.
There are several reasons for drawing the intractability line in this way. First,
the entries of Table 2.1 and their discussion in Section 2.1 imply that we cannot
solve arbitrary instances of intractable problems in a reasonable amount of time
394
Limitations of Algorithm Power
unless such instances are very small. Second, although there might be a huge
difference between the running times in O(p(n)) for polynomials of drastically
different degrees, there are very few useful polynomial-time algorithms with the
degree of a polynomial higher than three. In addition, polynomials that bound
running times of algorithms do not usually involve extremely large coefficients.
Third, polynomial functions possess many convenient properties; in particular,
both the sum and composition of two polynomials are always polynomials too.
Fourth, the choice of this class has led to a development of an extensive theory
called computational complexity, which seeks to classify problems according to
their inherent difficulty. And according to this theory, a problem's intractability
remains the same for all principal models of computations and all reasonable
input-encoding schemes for the problem under consideration.
We just touch on some basic notions and ideas of complexity theory in this
section. If you are interested in a more formal treatment of this theory, you will
have no trouble finding a wealth of textbooks devoted to the subject (e.g., [Sip97],
[Pap94]).
P and NP Problems
Most problems discussed in this book can be solved in polynomial time by some
algorithm. They include computing the product and the greatest common divisor
of two integers, sorting, searching (for a particular key in a list or for a given pattern
in a text string), checking connectivity and acyclicity of a graph, finding a minimum
spanning tree, and finding the shortest paths in a weighted graph. (You are invited
to add more examples to this list.) Informally, we can think about problems that
can be solved in polynomial time as the set that computer science theoreticians
call P. A more formal defmition includes in P only decision problems, which are
problems with yes/no answers.
DEFINITION 2 Class P is a class of decision problems that can be solved in
polynomial time by (deterministic) algorithms. This class of problems is called
polynomial.
The restriction of P to decision problems can be justified by the following
reasons. First, it is sensible to exclude problems not solvable in polynomial time
because of their exponentially large output. Such problems do arise naturallye.g., generating subsets of a given set or all the permutations of n distinct itemsbut it is apparent from the outset that they cannot be solved in polynomial time.
Second, many important problems that are not decision problems in their most
natural formulation can be reduced to a series of decision problems that are easier
to study. For example, instead of asking about the minimum number of colors
needed to color the vertices of a graph so that no two adjacent vertices are colored
the same color, we can ask whether there exists such a coloring of the graph's
vertices with no more than m colors for m = 1, 2, .... (The latter is called the
11.3 P, NP, and NP-complete Problems
395
m-coloring problem.) The first value of m in this series for which the decision
problem of m-coloring has a solution solves the optimization version of the graphcoloring problem as well.
It is natural to wonder whether every decision problem can be solved in
polynomial time. The answer to this question turns out to be no. In fact, some
decision problems cannot be solved at all by any algorithm. Such problems are
called undecidable. A famous example was given by Alan Turing in 19361 The
problem in question is called the halting problem: given a computer program and
an input to it, determine whether the program will halt on that input or continue
working indefinitely on it.
Here is a surprisingly short proof of this remarkable fact. By way of contradiction, assume that A is an algorithm that solves the halting problem. That is, for
any program P and input I,
A(P, I)= {
program P halts on input I;
~: ifif program
P does not halt on input I.
We can consider program P as an input to itself and use the output of algorithm
A for pair (P, P) to construct a program Q as follows:
Q(P) _ { halts, if A(P, P) = 0, i.e., if program P does not halt on input P;
- does not halt, if A(P, P) = 1, i.e., if program P halts on input P.
Then on substituting Q for P, we obtain
Q(Q) _ { halts, if A(Q, Q) = 0, i.e., if program Q does not halt on input Q;
- does not halt, if A(Q, Q) = 1, i.e., if program Q halts on input Q.
This is a contradiction because neither of the two outcomes for program Q is
possible, and it completes the proof.
Are there decidable but intractable problems? Yes, there are, but the number
of known examples is relatively small, especially of those that arise naturally rather
than being constructed for the sake of a theoretical argument.
There are many important problems, however, for which no polynomial-time
algorithm has been found, nor has the impossibility of such an algorithm been
proved. The classic monograph by M. Garey and D. Johnson [Gar79] contains a
list of several hundred such problems from different areas of computer science,
mathematics, and operations research. Here is just a small sample of some of the
best-known problems that fall into this category:
1.
This was just one of many breakthrough contributions to theoretical computer science made by the
English logician and computer science pioneer Alan Turing (1912-1954). In recognition of this, the
ACM-the principal society of computing professionals and researchers-has named after him an
award given to researchers for outstanding contributions to theoretical computer science. A lecture
given on such an occasion by Richard Karp [Kar86] provides an interesting historical account of the
development of complexity theory.
396
Limitations of Algorithm Power
Hamiltonian circuit Determine whether a given graph has a Hamiltonian
circuit (a path that starts and ends at the same vertex and passes through all
the other vertices exactly once).
Traveling salesman Find the shortest tour through n cities with known positive integer distances between them (find the shortest Hamiltonian circuit in
a complete graph with positive integer weights).
Knapsack problem Find the most valuable subset of n items of given positive
integer weights and values that fit into a knapsack of a given positive integer
capacity.
Partition problem Given n positive integers, determine whether it is possible
to partition them into two disjoint subsets with the same sum.
Bin packing Given n items whose sizes are positive rational numbers not
larger than 1, put them into the smallest number of bins of size 1.
Graph coloring For a given graph, find its chromatic number (the smallest
number of colors that need to be assigned to the graph's vertices so that no
two adjacent vertices are assigned the same color).
Integer linear programming Find the maximum (or minimum) value of a
linear function of several integer-valued variables subject to a finite set of
constraints in the form of linear equalities and/or inequalities.
Some of these problems are decision problems. Those that are not have
decision-version counterparts (e.g., them-coloring problem for the graph-coloring
problem). What all these problems have in common is an exponential (or worse)
growth of choices, as a function of input size, from which a solution needs to be
found. Note, however, that some problems that also fall under this umbrella can
be solved in polynomial time. For example, the Eulerian circuit problem-the
problem of the existence of a cycle that traverses all the edges of a given graph
exactly once-can be solved in O(n 2 ) time by checking, in addition to the graph's
connectivity, whether all the graph's vertices have even degrees. This example is
particularly striking: it is quite counterintuitive to expect that the problem about
cycles traversing all the edges exactly once (Eulerian circuits) can be so much
easier than the seemingly similar problem about cycles visiting all the vertices
exactly once (Hamiltonian circuits).
Another common feature of a vast majority of decision problems is the
fact that while solving such problems can be computationally difficult, checking
whether a proposed solution actually solves the problem is computationally easy,
i.e., it can be done in polynomial time. (We can think of such a proposed solution
as being randomly generated by somebody leaving us with the task of verifying
its validity.) For example, it is easy to check whether a proposed list of vertices
is a Hamiltonian circuit for a given graph with n vertices. All we need to check
is that the list contains n + 1 vertices of the graph in question, the first n vertices
are distinct while the last one is the same as the first, and every consecutive pair
of the list's vertices is connected by an edge. This general observation about de-
11.3
P, NP, and NP-complete Problems
397
cision problems has led computer scientists to the notion of a nondeterministic
algorithm.
DEFINITION 3 A nondeterministic algorithm is a two-stage procedure that
takes as its input an instance I of a decision problem and does the following.
Nondeterministic ("guessing") stage: An arbitrary string Sis generated that
can be thought of as a candidate solution to the given instance I (but may be
complete gibberish as well).
Deterministic ("verification") stage: A deterministic algorithm takes both I
and S as its input and outputs yes if S represents a solution to instance I. (If Sis
not a solution to instance I, the algorithm either returns no or is allowed not to
halt at alL)
We say that a nondeterministic algorithm solves a decision problem if and
only if for every yes instance of the problem it returns yes on some execution. (In
other words, we require a nondeterministic algorithm to be capable of "guessing"
a solution at least once and to be able to verify its validity. And, of course, we
do not want it to ever output a yes answer on an instance for which the answer
should be no.) Finally, a nondeterministic algorithm is said to be nondeterministic polynomial if the time efficiency of its verification stage is polynomiaL
Now we can define the class of NP problems.
DEFINITION 4 Class NP is the class of decision problems that can be solved by
nondeterministic polynomial algorithms. This class of problems is called nondeterministic polynomial.
Most decision problems are in NP. First of all, this class includes all the
problems in P:
P s; NP.
This is true because, if a problem is in P, we can use the deterministic polynomialtime algorithm that solves it in the verification-stage of a nondeterministic algorithm that simply ignores string S generated in its nondeterministic ("guessing")
stage. But N P also contains the Hamiltonian circuit problem, the partition problem, as well as decision versions of the traveling salesman, the knapsack, graph
coloring and many hundreds of other difficult combinatorial optimization problems cataloged in [Gar79]. The halting problem, on the other hand, is among the
rare examples of decision problems that are known not to be in NP.
This leads to the most important open question of theoretical computer science: Is P a proper subset of NP, or are these two classes, in fact, the same? We
can put this symbolically as
'I
P~NP.
y
398
Limitations of Algorithm Power
Note that P = NP would imply that each of many hundreds of difficult combinatorial decision problems can be solved by a polynomial-time algorithm, although computer scientists have failed to find such algorithms despite their persistent efforts over many years. Moreover, many well-known decision problems
are known to be "NP-complete" (see below), which seems to cast more doubts
on the possibility that P = N P.
NP-Complete Problems
Informally, an NP-complete problem is a problem in NP that is as difficult as any
other problem in this class because, by definition, any other problem in N P can
be reduced to it in polynomial time (shown symbolically in Figure 11.6).
Here are more formal definitions of these concepts.
DEFINITION 5 A decision problem D1 is said to be polynomially reducible to
a decision problem D 2 if there exists a function t that transforms instances of D1
to instances of D2 such that
1.
t maps all yes instances of D 1 to yes instances of D 2 and all no instances of Dt
to no instances of D 2;
2.
tis computable by a polynomial-time algorithm.
This definition immediately implies that if a problem D 1 is polynomially
reducible to some problem D 2 that can be solved in polynomial time, then problem
D1 can also be solved in polynomial time (why?).
NPproblems
·r;
NP-complete
/
FIGURE 11.6 Notion of an NP-complete problem. Polynomial-time reductions of NP
problems to an NP-complete problem are shown by arrows.
11.3
DEFINITION 6
1.
2.
P, NP, and NP-complete Problems
399
A decision problem Dis said to be NP-complete if
it belongs to class NP;
every problem in NP is polynomially reducible to D,
The fact that closely related decision problems are polynomially reducible to
each other is not very surprising. For example, let us prove that the Hamiltonian
circuit problem is polynomially reducible to the decision version of the traveling
salesman problem. The latter can be stated as the problem to determine whether
there exists a Hamiltonian circuit in a given complete graph with positive integer
weights whose length is not greater than a given positive integer m. We can map
a graph G of a given instance of the Hamiltonian circuit problem to a complete
weighted graph G' representing an instance of the traveling salesman problem by
assigning 1 as the weight to each edge in G and adding an edge of weight 2 between
any pair of nonadjacent vertices in G, As the upper bound m on the Hamiltonian
circuit length, we take m = n, where n is the number of vertices in G (and G').
Obviously, this transformation can be done in polynomial time.
Let G be a yes instance of the Hamiltonian circuit problem. Then G has a
Hamiltonian circuit, and its image in G' will have length n, making the image a
yes instance of the decision traveling salesman problem. Conversely, if we have
a Hamiltonian circuit of the length not larger than n in G', then its length must
be exactly n (why?) and hence the circuit must be made up of edges presentinG,
making the inverse image of the yes instance of the decision traveling salesman
problem be a yes instance of the Hamiltonian circuit problem. This completes the
proof.
The notion of NP-completeness requires, however, polynomial reducibility
of all problems in NP, both known and unknown, to the problem in question.
Given the bewildering variety of decision problems, it is nothing short of amazing
that the specific examples of NP-complete problems have been actually found.
Nevertheless, this mathematical feat was accomplished independently by Stephen
Cook in the United States and Leonid Levin in the former Soviet Union 2 In his
1971 paper, Cook [Coo71] showed that the so-called CNF-satisfiability pmblem
is NP-complete. The CNF-satisfiability problem deals with boolean expressions.
Each boolean expression can be represented in conjunctive normal form, such as
the following expression involving three boolean variables x 1, x 2 , and x 3 and their
negations denoted il> x2 , and x3 , respectively:
2.
As it often happens in the history of science, breakthrough discoveries arc made independently and
almost simultaneously by several scientists. In fact, Levin introduced a more general notion than NPcompleteness, which was not limited to decision problems, but his paper rLev73] was published two
years after Cook's.
400
Limitations of Algorithm Power
The CNF-satisfiability problem asks whether or not we can assign values true and
false to variables of a given boolean expression in its CNF form to make the entire
expression true. (It is easy to see that this can be done for the above formula: if
x1 =true, x2 =true, and x3 =false, the entire expression is true.)
Since the Cook-Levin discovery of the first known NP-complete problems,
computer scientists have found many hundreds, if not thousands, of other examples. In particular, the well-known problems (or their decision versions) mentioned above-Hamiltonian circuit, traveling salesman, partition, bin packing,
and graph coloring-are all NP-complete. It is known, however, that if P ,P NP,
there must exist NP problems that neither are in P nor are NP-complete.
For a while, the leading candidate to be such an example was the problem
of determining whether a given integer is prime or composite. But in an important theoretical breakthrough, Professor Manindra Agrawal and his students
Neeraj Kayal and Nitin Saxena of the Indian Institute of Technology in Kanpur
announced in 2002 the discovery of a deterministic polynomial-time algorithm
for primality testing (Agr02]. Their algorithm does not solve, however, therelated problem of factoring large composite integers, which lies at the heart of the
widely used encryption method called the RSA algorithm [Riv78].
Showing that a decision problem is NP-complete can be done in two steps.
First, one needs to show that the problem in question is in NP; i.e., a randomly
generated string can be checked in polynomial time to determine whether or not
it represents a solution to the problem. Typically, this step is easy. The second
step is to show that every problem in NP is reducible to the problem in question
in polynomial time. Because of the transitivity of polynomial reduction, this step
can be done by showing that a known NP-complete problem can be transformed
to the problem in question in polynomial time (see Figure 11.7). Although such
a transformation may need to be quite ingenious, it is incomparably simpler than
proving the existence of a transformation for every problem in NP. For example,
if we already know that the Hamiltonian circuit problem is NP-complete, its
polynomial reducibility to the decision traveling salesman problem implies that
the latter is also NP-complete (after an easy check that the decision traveling
salesman problem is in class NP).
The definition of NP-completeness immediately implies that if there exists a
deterministic polynomial-time algorithm for just one NP-complete problem, then
every problem in N P can be solved in polynomial time by a deterministic algorithm, and hence P = NP. In other words, finding a polynomial-time algorithm
for one NP-complete problem would mean that there is no qualitative difference
between the complexity of checking a proposed solution and finding it in polynomial time for the vast majority of decision problems of all kinds. Such implications
make most computer scientists believe that P ,P NP, although nobody has been
successful so far in finding a mathematical proof of this intriguing conjecture. Surprisingly, in interviews with the authors of a book about the lives and discoveries
of 15 prominent computer scientists (Sha98], Cook seemed to be uncertain about
11.3
P, NP, and NP-complete Problems
401
NP problems
known
NP-complete
problem
candidate for
·f'~-----.._~ NP-completeness
FIGURE 11.7 Proving NP-completeness by reduction
the eventual resolution of this dilemma while Levin contended that we should
expect the P = NP outcome.
=
?
Whatever the eventual answer to the P NP question proves to be, knowing
that a problem is NP-complete has important practical implications for today. It
means that faced with a problem known to be NP-complete, we should probably
not aim at gaining fame and fortune 3 by designing a polynomial-time algorithm
for solving all its instances. Rather, we should concentrate on several approaches
that seek to alleviate the intractability of such problems. These approaches are
outlined in the next chapter of the book.
-----Exercises 1 1 . 3 - - - - - - - - - - - - - - 1. A game of chess can be posed as the following decision problem: given a
legal positioning of chess pieces and information about which side is to move,
determine whether that side can win. Is this decision problem decidable?
2. A certain problem can be solved by an algorithm whose running time is in
O(nlog, ").Which of the following assertions is true?
a. The problem is tractable.
b. The problem is intractable.
c. None of the above.
3.
In 2000, The Clay Mathematics Institute (CMI) of Cambridge, Massachusetts, designated a $1 million
prize for the solution to this problem.
402
Limitations of Algorithm Power
3. Give examples of the following graphs or explain why such examples cannot
exist.
a. graph with a Hamiltonian circuit but without an Eulerian circuit
b. graph with an Eulerian circuit but without a Hamiltonian circuit
c. graph with both a Hamiltonian circuit and an Eulerian circuit
d. graph with a cycle that includes all the vertices but with neither a Hamiltonian circuit nor an Eulerian circuit
4. For each of the following graphs, find its chromatic number.
a.
b.
5. Design a polynomial-time algorithm for the graph 2-coloring problem: determine whether vertices of a given graph can be colored in no more than two
colors so that no two adjacent vertices are colored the same color.
6. Consider the following brute-force algorithm for solving the composite number problem: Check successive integers from 2 to Ln/2J as possible divisors of
n. If one of them divides n evenly, return yes (i.e., the number is composite),
if none of them does, return no. Why does this algorithm not put the problem
in class P?
7. State the decision version for each of the following problems and outline a
polynomial-time algorithm that verifies whether or not a proposed solution
solves the problem. (You may assume that a proposed solution represents a
legitimate input to your verification algorithm.)
a. knapsack problem
b. bin-packing problem
8. Show that the partition problem is polynomially reducible to the decision
version of the knapsack problem.
T
I
11.3 P, NP, and NP-complete Problems
403
9. Show that the following three problems are polynomially reducible to each
other.
Determine, for a given graph G = (V, E) and a positive integer m ~ lVI,
whether G contains a clique of size m or more. (A clique of size kin a graph
is its complete subgraph of k vertices.)
Determine, for a given graph G = (V, E) and a positive integer m ~/VI,
whether there is a vertex cover of size m or less for G. (A vertex cover of
size k for a graph G = (V, E) is a subset V' <; V such that IV'I = k and, for
each edge (u, v) E E, at least one of u and v belongs to V'.)
Determine, for a given graph G = (V, E) and a positive integer m ~ /VI,
whether G contains an independent set of size m or more. (An independent
set of size k for a graph G = (V, E) is a subset V' <; V such that IV' I= k and
for all u, v E V', vertices u and v are not adjacent in G.)
10. Which of the following diagrams do not contradict the current state of our
knowledge about the complexity classes P, NP, and NPC (NP-complete
problems)?
a.
b.
P= NP= NPC
c.
d.
NP
P
NP
NPC
e.
NP
I
1.
~- _.·
..
..••·•.··.
.•.
Jl
Limitations of Algorithm Power
404
11. King Arthur expects 150 knights for an annual dinner at Camelot. Unfortunately, some of the knights quarrel with each other, and Arthur knows who
quarrels with whom. Arthur wants to seat his guests around a table so that no
two quarreling knights sit next to each other.
a. Which standard problem can be used to model King Arthur's task?
b. As a research project, find a proof that Arthur's problem has a solution if
each knight does not quarrel with at least 75 other knights.
11.4 Challenges of Numerical Algorithms
Numerical analysis is usually described as the branch of computer science concerned with algorithms for solving mathematical problems. This description needs
an important clarification: the problems in question are problems of "continuous"
mathematics-solving equations and systems of equations, evaluating such functions as sin x and In x, computing integrals, and so on-as opposed to problems of
discrete mathematics dealing with such structures as graphs, trees, permutations,
and combinations. Our interest in efficient algorithms for mathematical problems
stems from the fact that these problems arise as models of many real-life phenomena both in the natural world and in the social sciences. In fact, numerical
analysis used to be the main area of research, study, and application of computer
science. With the rapid proliferation of computers in business and everyday-life
applications, which deal primarily with storage and retrieval of information, the
relative importance of numerical analysis has shrunk in the last 30 years. However,
its applications, enhanced by the power of modern computers, continue to expand
in all areas of fundamental research and technology. Thus, wherever one's interests lie in the wide world of modern computing, it is important to have at least
some understanding of the special challenges posed by continuous mathematical
problems.
We are not going to discuss the variety of difficulties posed by modeling, the
task of describing a real-life phenomenon in mathematical terms. Assuming that
this has already been done, what principal obstacles to solving a mathematical
problem do we face? The first major obstacle is the fact that most numerical analysis problems cannot be solved exactly4 They have to be solved approximately, and
this is usually done by replacing an infinite object by a finite approximation. For
example, the value of ex at a given point x can be computed by approximating
its infinite Taylor's series about x = 0 by a finite sum of its first terms, called the
nth-degree Taylor polynomial:
x2
xn
2!
n!
ex""l+x+-+···+-.
4.
(11.6)
Solving a system of linear equations and polynomial evaluation discussed in Sections 6.2 and 6.5,
respectively, are rare exceptions to this rule.
~
I
11.4
Challenges of Numerical Algorithms
405
l
--~r-~h~4-----------~h__._~h__.___________._~h__._____.x
a
X;
Xn-1
b
FIGURE 11.8 Composite trapezoidal rule
To give another example, the definite integral of a function can be approximated
by a finite weighted sum of its values, as in the composite tmpezoidal rule that
you might remember from your calculus class:
1
J
b
t;
n-1
f(x)dx"' i[f(a) + 2
f(x 1) + f(h)],
(11.7)
where h = (b- a)jn, x; =a+ ih fori= 0, 1, ... , n (Figure 11.8).
The errors of such approximations are called (J'Uncation errors. One of the
major tasks in numerical analysis is to estimate the magnitudes of truncation
errors. This is typically done by using calculus tools, from elementary to quite
advanced. For example, for approximation (11.6) we have
2
II
M
lex- [1 + x + :'".... + · · · + :'"...]1 S
lxl"+l,
2!
n!
(n+1)!
(11.8)
where M =max e~ on the segment with the endpoints at 0 and x. This formula
makes it possible to determine the degree of Taylor's polynomial needed to guarantee a predefined accuracy level of approximation (11.6).
For example, if we want to compute e0 ·5 by formula (11.6) and guarantee the
truncation error to be smaller than w- 4 , we can proceed as follows. First, we
estimate M of formula (11.8):
M = max eS < e 0·5 < 2.
O::s{:s0.5
-
Using this bound and the desired accuracy level of 10-4 , we obtain from (11.8)
M
(n+1)!
10.51"+ 1 <
2
(n+l)!
os'+l < w- 4
n
..
····I·
406
Limitations of Algorithm Power
To solve the last inequality, we can compute the first few values of
I
2
2- 11
c-----c--:0.5''+1 = - - (n + 1)!
(n + 1)!
to see that the smallest value of n for which this inequality holds is 5.
Similarly, for approximation (11.7), the standard bound of the truncation
error is given by the inequality
b
I
1
h
f(x)dx- -[f(a)
a
2
n-1
+ 2 L f(x;) + f(b)]l S:
;~J
(b _
)h2
a
12
M2,
(11.9)
where M 2 =max lf"(x)! on the interval a::: x::: b. You are asked to use this
inequality in the exercises for this section (Problems 5 and 6).
The other type of errors, called round-off errors, are caused by the limited
accuracy with which we can represent real numbers in a digital computer. These
errors arise not only for all irrational numbers (which, by definition, require an
infinite number of digits for their exact representation) but for many rational
numbers as well. In the overwhelming majority of situations, real numbers are
represented as floating-point numbers,
(11.10)
where B is the number base, usually 2 or 16 (or, for unsophisticated calculators,
10); d 1, d2 , ... , dP are digits (0::: d; < B fori= 1, 2, ... , p and d1 > 0 unless the
number is 0) representing together the fractional part of the number and called
its mantissa; E is an integer exponent with the range of values approximately
symmetric about 0.
The accuracy of the floating-point representation depends on the number
of significant digits p in representation (11.10). Most computers permit two or
even three levels of precision: single precision (typically equivalent to between
6 and 7 significant decimal digits), double precision (13 to 14 significant decimal
digits), and extended precision (19 to 20 significant decimal digits). Using higherprecision arithmetic slows computations but may help to overcome some of the
problems caused by round-off errors. Higher precision may need to be used only
for a particular step of the algorithm in question.
As with an approximation of any kind, it is important to distinguish between
the absolute error and the relative error of representing a number a* by its
approximation a:
absolute error= Ia -a* I.
relative error =
(The relative error is undefined if a*= 0.)
Ia- a* I
Ia*!
.
(11.11)
(11.12)
11.4
Challenges of Numerical Algorithms
407
Very large and very small numbers cannot be represented in floating-poiut
arithmetic because of the phenomena called oveljlow and undetjlow, respectively. An overflow happens when an arithmetic operation yields a result outside the range of the computer's floating-point numbers. Typical examples of
overflow arise from the multiplication of large numbers or division by a very
small number. Sometimes we can eliminate this problem by making a simple
change in the order in which an expression is evaluated (e.g., (1029. 1130)/1230 =
29
30
10 . (11/12) ), by replacing an expression with an equal one (e.g., computing
1 0
( g) not as 100!/(2!(100- 2)!) but as (100 · 99)/2), or by computing a logarithm
of an expression instead of the expression itself.
Underflow occurs when the result of an operation is a nonzero fraction of
such a small magnitude that it cannot be represented as a nonzero floating-point
number. Usually, underflow numbers are replaced by zero, but a special signal is
generated by hardware to indicate such an event has occurred.
It is important to remember that, in addition to inaccurate representation
of numbers, the arithmetic operations performed in a computer are not always
exact, either. In particular, subtracting two nearly equal floating-point numbers
may cause a large increase in relative error. This phenomenon is called subtractive
cancellation.
EXAMPLE 1 Consider two irrational numbers
a*="= 3.14159265 ... and fl* ="- 6. 10"" 7 = 3.14159205 ...
represented by floating-point numbers a= 0.3141593. 101 and f1 = 0.3141592.
101, respectively. The relative errors of these approximations are small:
Ia- a* I = 0.0000003 ... < ~
a*
n
3
10
_7
and
If!- fl*l = o.ooooooo5 ... < !
fl*
n-6-10-7
_
10 7
3'
respectively. The relative error of representing the difference y' =a* - fl* by the
difference of the floating-point representations y = a - f1 is
IY- y*l
w- 6 - 6 . 10-7
2
y*
6-10-7
3'
which is very large for a relative error despite quite accurate approximations for
E
~a~fl.
Note that we may get a significant magnification of round-off error if such a
low-accuracy difference is used as a divisor. (We already encountered this problem
in discussing Gaussian elimination in Section 6.2. Our solution there was to use
partial pivoting.) Many numerical algorithms involve thousands or even millions
408
Limitations of Algorithm Power
of arithmetic operations for typical inputs. For such algorithms, the propagation of
round-off errors becomes a major concern from both the practical and theoretical
standpoints. For some algorithms, round-off errors can propagate through the
algorithm's operations with increasing effect. This highly undesirable property
of a numerical algorithm is called instability. Some problems exhibit such a high
level of sensitivity to changes in their input that it is all but impossible to design a
stable algorithm to solve them. Such problems are called ill-conditioned.
EXAMPLE 2 Consider the following system of two linear equations in two
unknowns:
1.00lx + 0.999y = 2
0.999x + 1.001y = 2.
Its only solution is x = 1, y = 1. To see how sensitive this system is to small changes
to its tight-hand side, consider the system with the same coefficient matrix hut
slightly different right-hand side values:
l.OOlx + 0.999y = 2.002
0.999x + l.OOly = 1.998.
The only solution to this system is x = 2, y = 0, which is quite far from the solution
to the previous system. Note that the coefficient matrix of this system is close to
being singular (why?). Hence, a minor change in its coefficients may yield a system with either no solutions or infinitely many solutions, depending on its righthand side values. You can find a more formal and detailed discussion of how we can
measure the degree of ill-condition of the coefficient matrix in numerical analysis
textbooks (e.g., [Ger03]).
1!1i
We conclude with a well-known problem of finding real roots of the quadratic
equation
(11.13)
for any real coefficients a, b, and c (a oJ 0). According to secondary-school algebra,
2
equation (11.13) has real roots if and only if its discriminant D = b - 4ac is
nonnegative, and these roots can be found by the following formula:
Xt,2 =
-b± )b2 - 4ac
2a
(11.14)
Although formula (11.14) provides a complete solution to the posed problem
as far as a mathematician is concerned, it is far from being a complete solution
for an algorithm designer. The first major obstacle is evaluating the square root.
Even for most integers D, .fl5 is an irrational number that can be computed only
approximately. There is a method of computing square roots that is much better
than the one commonly taught in secondary school. (It follows from Newton's
11.4
Challenges of Numerical Algorithms
409
method, a very important algorithm for solving equations, which we discuss in
Section 12.4.). This method generates the sequence {x,] of approximations to v'D,
where D is a given nonnegative number, according to the formula
1
2
D
x,+l = - (x, + -) for n = 0, 1, ... ,
X 11
(11.15)
where the initial approximation x 0 can be chosen (among other possibilities) as
x 0 = (1 + D)/2. It is not difficult to prove (Problem 9a) that sequence (11.15) is
strictly decreasing (if D oft1) and converges to v'D. We can stop generating its
elements either when the difference between its two consecutive elements is less
than a predefined error tolerance E > 0
Xn- Xn+l < E
or when x,;+l is sufficiently close to D. Approximation sequence (11.15) converges
very fast to v'D for most values of D. In particular, one can prove that if 0.25 :<=
D < 1, then no more than four iterations are needed to guarantee that
lx,- vfril < 4. w- 15 ,
and we can always scale a given value of d to one in the interval [0.25, 1) by the
formula d = D2P, where p is an even integer.
EXAMPLE 3 Let us apply Newton's algorithm to compute ,/2. (For simplicity,
we ignore sealing.) We round off the numbers to six decimal places and use the
standard numerical analysis notation ""' to indicate the round-offs.
xo =
1
2(1 + 2) = 1:500000
1
2
2
x0
x 1 = - (x 0 + -) ""'1.416667
1
2
2
Xj
1
2
x 3 = -(x2 + -) ""'1.414214
2
x2
1
2
x4 = -(x3 + -) ""'1.414214.
2
X3
Xz = -(Xj + - ) ""'1.414216
At this point we have to stop because x 4 = x 3 ""' 1.414214, and hence all other
approximations will be the same. The exact value of ,/2 is 1.41421356....
1111
With the issue of computing square roots squared away (I do not know
whether or not the pun was intended), are we home free to write a program based
Limitations of Algorithm Power
410
on formula (11.14)? The answer is no because of the possible impact of round-off
errors. Among other obstacles, we are faced here with the menace of subtractive
cancellation. If b2 is much larger than4ac, ,/b2- 4ac will be very close to Ihi, and
a root computed by formula (11.14) might have a large relative error.
EXAMPLE 4 Let us follow a paper by George Forsythe5 [For69] and consider
the equation
Its true roots to 11 significant digits are
x{ ""99999.999990
and
x; ""0.000010000000001.
If we use formula (11.14) and perform all the computations in decimal floating-
point arithmetic with, say, seven significant digits, we obtain
( -b) 2 = 0.1000000. 1011
4ac = 0.4000000 . Jo1
D ""0.1000000. 1011
FD"" o.1oooooo. 106
X]"" -b
+ FD ""0.1000000. 106
2a
. -h- .fD .
x2 =
=0.
2a
And while the relative error of approximating x{ by x1 is very small, for the second
root it is very large:
lx2- x21
x2
~-""- = 1
(i.e., 100%).
To avoid the possibility of substractive cancellation in formula (11.14 ), we can
use instead another formula, obtained as follows:
5.
George E. Forsythe (1917~1972), a noted numerical analyst, played a leading role in establishing
computer science as a separate academic discipline in the United States. It is his words that are used
as the epigraph to this book's preface.
11.4
Challenges of Numerical Algorithms
411
2a
( -b + ,Jb2- 4ac)( -b- .jb2- 4ac)
2a( -b- ,Jb 2 - 4ac)
=--.::;o2c~=
-b - ,Jb 2 - 4ac'
with no danger of subtractive cancellation in the denominator if b > 0. As to x ,
2
it can be computed by the standard formula
2a
with no danger of cancellation, either, for a positive value of b.
The case of b < 0 is symmetric: we can use the formulas
2a
and
2c
x2 = ---'~===
-b + ,Jb2- 4ac
(The case of b = 0 can be considered with either of the other two cases.)
There are several other obstacles to applying formula (11.14), which are related to limitations of floating-point arithmetic: if a is very small, division by a
can cause an overflow; there seems to be no way to fight the danger of subtractive
2
cancellation in computing b - 4ac other than calculating it with double precision;
and so on. These problems have been overcome by William Kahan of the University of Toronto (see [For69]), and his algorithm is considered to be a significant
achievement in the history of numerical analysis.
Hopefully, this brief overview has piqued your interest enough for you to seek
more information in the many books devoted exclusively to numerical algorithms.
In this book, we discuss one more topic in the next chapter: three classic methods
for solving equations in one unknown.
-----Exercises 1 1 . 4 - - - - - - - - - - - - - - 1. Some textbooks define the number of significant digits in the approximation
of number a* by number a as the largest nonnegative integer k for which
/a -a*/ < 5. 10-k
/a*/
412
Limitations of Algorithm Power
According to this definition, how many significant digits are there in the
approximation of" by
a. 3.1415?
b. 3.1417?
2. If a = 1.5 is known to approximate some number a' with the absolute error
not exceeding 10-2 , find
a. the range of possible values of a'.
b. the range of the relative errors of these approximations.
3. Find the approximate valne of .je = 1.648721 ... obtained by the fifth-degree
Taylor's polynomial about 0 and compute the truncation error of this approximation. Does the result agree with the theoretical prediction made in the
section?
4. Derive formula (11.7) of the composite trapezoidal rule.
5. Use the composite trapezoidal rule with n = 4 to approximate the following
definite integrals. Find the truncation error of each approximation and compare it with the one given by formula (11.9).
rl X zG1X
a. JO
b • !3
l X -ld X
6. If J~ esin x dx is to be computed by the composite trapezoidal rule, how large
should the number of subintervals be to guarantee a truncation error smaller
than 10-4 ? What about 10-6 ?
7. Solve the two systems of linear equations and indicate whether they are illconditioned.
2x +5y =7
a. 2x + 5.000001y = 7.000001
2x +5y =7
b. 2x + 4.999999y = 7.000002
8. Write a computer program for solving the equation ax 2 + bx + c = 0.
9. a. Prove that for any nonnegative number D, the sequence of Newton's
method for computing ~ is strictly decreasing and converges to ~for
any value of the initial approximation x 0 > ~b. Prove that if 0.25 :S D < 1 andx0 = (1 + D)/2,nomore than four iterations
of Newton's method are needed to guarantee that
lx,- ~I < 4 · 10- 15
10. Apply four iterations of Newton's method to compute ~ and estimate the
absolute and relative errors of this approximation.
,,,,~,,,,,
I
Summary
413
SUMMARY
"
Given a class of algorithms for solving a particular problem, a lower bound
indicates the best possible efficiency any algorithm from this class can have.
"
A trivial lower bound is based on counting the number of items in the
problem's input that must be processed and the number of output items
that need to be produced.
"
An information-theoretic lower bound is usually obtained through a mechanism of decision trees. This technique is particularly useful for comparisonbased algorithms for sorting and searching. Specifically:
Any general comparison-based sorting algorithm must perform at least
flog 2 n !l ""n 1og2 n key comparisons in the worst case.
Any general comparison-based algorithm for searching a sorted array
must perform at least flog 2 (n + 1) l key comparisons in the worst case.
II
The adversary method for establishing lower bounds is based on following
the logic of a malevolent adversary who pushes the algorithm into the most
time-consuming path.
"
A lower bound can also be established by reduction, i.e., by reducing a
problem with a known lower bound to the problem in question.
"
Complexity theory seeks to classify problems according to their computational
complexity. The principal split is between tractable and intractable problemsproblems that can and cannot be solved in polynomial time, respectively.
For purely technical reasons, complexitY, theory concentrates on decision
problems, which are problems with yes/no answers.
,.
The halting problem is an example of an undecidable decision problem; i.e.,
it cannot be solved by any algorithm.
"
P is the class of all decision problems that can be solved in polynomial time.
NP is the class of all decision problems whose randomly guessed solutions
can be verified in polynomial time.
"
Many important problems in NP (such as the Hamiltonian circuit problem)
are known to be NP -complete: all other problems in NP are reducible to such
a problem in polynomial time. The first proof of a problem's NP-completeness
was published by S. Cook for the CNF-satisfiability problem.
"
It is not known whether P = NP or P is just a proper subset of NP. This
question is the most important unresolved issue in theoretical computer
science. A discovery of a polynomial-time algorithm for any of the thousands
known NP-complete problems would imply that P = NP.
----,
!
414
Limitations of Algorithm Power
"'
Numerical analysis is a branch of computer science dealing with solving
continuous mathematical problems. Two types of errors occur in solving a
majority of such problems: truncation error and round-off error. Truncation
errors stem from replacing infinite objects by their finite approximations.
Round-off errors are due to inaccuracies of representing numbers in a digital
computer.
"'
Subtractive cancellation happens as a result of subtracting two near-equal
floating-point numbers. It may lead to a sharp increase in the relative roundoff error and therefore should be avoided (by either changing the expression's
form or by using a higher precision in computing such a difference).
"'
Writing a general computer program for solving quadratic equations ax 2 +
bx + c = 0 is a difficult task. The problem of computing square roots can be
solved byutilitizing Newton's method; the problem of subtractive cancellation
can be dealt with by using different formulas depending on whether coefficient
b is positive or negative and by computing the discriminant b2 - 4ac with
double precision.
Coping with the
limitations of
Algorithm Power
Keep on the lookout for novel ideas that others have used successfully.
Your idea has to be original only in its adaptation to the problem you're
working on.
-Thomas Edison (1847-1931)
A
s we saw ill the previous chapter, there are problems that are difficult to
solve algorithmically. At the same time, some of them are so important that
we cannot just sigh in resignation and do nothing. This chapter outlines several
ways of dealing with such difficult problems.
Sections 12.1 and 12.2 introduce two algorithm design techniques-backtracking and branch-and-bound-that often make it possible to solve at least
some large instances of difficult combinatorial problems. Both strategies can be
considered an improvement over exhaustive search, discussed in Section 3.4.
Unlike exhaustive search, they construct candidate solutions one component at a
time and evaluate the partially constructed solutions: if no potential values of the
remaining components can lead to a solution, the remaining components are not
generated at all. This approach makes it possible to solve some large instances of
difficult combinatorial problems, though, in the worst case, we still face the same
curse of exponential explosion encountered in exhaustive search.
Both backtracking and branch-and-bound are based on the construction of a
state-space tree whose nodes reflect specific choices made for a solution's components. Both techniques terminate a node as soon as it can be guaranteed that no
solution to the problem can be obtained by considering choices that correspond
to the node's descendants. The techniques differ in the nature of problems they
can be applied to. Branch-and-bound is applicable only to optimization problems
because it is based on computing a bound on possible values of the problem's objective function. Backtracking is not constrained by this demand, but more often
415
I'
I
I
416
Coping with the Limitations of Algorithm Power
than not, it applies to nonoptimization problems. The other distinction between
backtracking and branch-and-bound lies iu the order in which nodes of the statespace tree are generated. For backtracking, this tree is usually developed depth
first (i.e., similar to DFS). Branch-and-bound can generate nodes according to
several rules; the most natural one is the so-called best -first rule explained in Section 12.2.
Section 12.3 takes a break from the idea of solving a problem exactly. The
algorithms presented there solve problems approximately but fast. Specifically,
we consider a few approximation algorithms for the traveling salesman and knapsack problems. For the traveling salesman problem, we discuss basic theoretical
results and pertinent empirical data for several well-known approximation algorithms. For the knapsack problem, we first introduce a greedy algorithm and then
a parametric family of polynomial-time algorithms that yield arbitrarily good approximations.
Section 12.4 is devoted to algorithms for solving nonlinear equations. After a
brief discussion of this very important problem, we examine three classic methods
for approximate root finding: the bisection method, the method of false position,
and Newton's method.
12.1
Backtracking
Throughout the book (see in particular Sections 3.4 and 11.3), we have encountered problems that require finding an element with a special property in a domain
that grows exponentially fast (or faster) with the size of the problem's input: a
Hamiltonian circuit among all permutations of a graph's vertices, the most valuable subset of items for an instance of the knapsack problem, and the like. We
addressed in Section 11.3 the reasons for believing that many such problems might
not be solvable in polynomial time. Also recall that we discussed in Section 3.4
how such problems can be solved, at least in principle, by exhaustive search. The
exhaustive-search technique suggests generating all candidate solutions and then
identifying the one (or the ones) with a desired property.
Backtracking is a more intelligent variation of this approach. The principal
idea is to construct solutions one component at a time and evaluate such partially
constructed candidates as follows. If a partially constructed solution can be developed further without violating the problem's constraints, it is done by taking
the first remaining legitimate option for the next component. If there is no legitimate option for the next component, no alternatives for any remaining component
need to be considered. In this case, the algorithm backtracks to replace the last
component of the partially constructed solution with its next option.
It is convenient to implement this kind of processing by constructing a tree
of choices being made, called the state-space tree. Its root represents an initial
state before the search for a solution begins. The nodes of the first level in the
tree represent the choices made for the first component of a solution, the nodes
of the second level represent the choices for the second component, and so
12.1
Backtracking
417
on. A node in a state-space tree is said to be promising if it corresponds to a
partially constructed solution that may still lead to a complete solution; otherwise,
it is called nonpmmising. Leaves represent either nonpromising dead ends or
complete solutions found by the algorithm. In the majority of cases, a statespace tree for a backtracking algorithm is constructed in the manner of depthfirst search. If the current node is promising, its child is generated by adding the
first remaining legitimate option for the next component of a solution, and the
processing moves to this child. If the current node turns out to be nonpromising,
the algorithm backtracks to the node's parent to consider the next possible option
for its last component; if there is no such option, it backtracks one more level up
the tree, and so on. Finally, if the algorithm reaches a complete solution to the
problem, it either stops (if just one solution is required) or continues searching
for other possible solutions.
n-Oueens Problem
As our first example, we use a perennial favorite of textbook writers, then-queens
problem. The problem is to place n queens on an n-by-n chessboard so that no two
queens attack each other by being in the same row or in the same column or on
the same diagonal. For n = 1, the problem has a trivial solution, and it is easy to
see that there is no solution for n = 2 and n = 3. So let us consider the four-queens
problem and solve it by the backtracking technique. Since each of the four queens
has to be placed in its own row, all we need to do is to assign a column for each
queen on the board presented in Figure 12.1.
We start with the empty board and then place queen 1 in the first possible
position of its row, which is in column 1 of row 1. Then we place queen 2, after
trying unsuccessfully colunms 1 and 2, in the first acceptable position for it, which
is square (2,3), the square in row 2 and column 3. This proves to be a dead end
because there is no acceptable position for queen 3. So, the algorithm backtracks
and puts queen 2 in the next possible position at (2,4). Then queen 3 is placed at
(3,2), which proves to be another dead end. The algorithm then backtracks all the
way to queen 1 and moves it to (1,2). Queen 2 then goes to (2,4), queen 3 to (3,1),
and queen 4 to (4,3), which is a solution to the problem. The state-space tree of
this search is shown in Figure 12.2.
2
3
4
l~
+-- queen 1
+-- queen 2
+-- queen 3
+-- queen 4
FIGURE 12.1 Board for the four-queens problem
Coping with the Limitations of Algorithm Power
418
0
5
1
Q
Q
~
2/
2
X
1
Q
Q
X
1
X
/l \ l/
X
X
X
Q
Q
X
4
Q
X
3
4
X
X
j \
X
X
7/
Q
Q
Q
X
X
X
Q
Q
Q
/j\
6
Q
4
X
/ i ~8
1
X
2
Q
X
Q
Q
Q
solution
FIGURE 12.2 State-space tree of solving the four-queens problem by backtracking.
x denotes an unsuccessful attempt to place a queen in the indicated
column. The numbers above the nodes indicate the order in which the
nodes are generated.
If other solutions need to be found (how many of them are there for the fourqueens problem?), the algorithm can simply resume its operations at the leaf at
which it stopped. Alternatively, we can use the board's symmetry for this purpose.
Hamiltonian Circuit Problem
As our next example, let us consider the problem of finding a Hamiltonian circuit
in the graph in Figure 12.3a.
12.1
Backtracking
419
0
a
1
b
2
9
c
6
e
7
d
8
dead end
dead end
f
dead end
solution
(a)
(b)
FIGURE 12.3 (a) Graph. (b) State-space tree for finding a Hamiltonian circuit. The
numbers above the nodes of the tree indicate the order in which the
nodes are generated.
Without loss of generality, we can assume that if a Hamiltonian circuit exists,
it starts at vertex a. Accordingly, we make vertex a the root of the state-space
tree (Figure 12.3b). The first component of our future solution, if it exists, is a first
intermediate vertex of a Hamiltonian circuit to be constructed. Using the alphabet
order to break the three-way tie among the vertices adjacent to a, we select vertex
b. From b, the algorithm proceeds to c, then to d, then toe, and finally to f, which
proves to be a dead end. So the algorithm backtracks from f toe, then to d, and
then to c, which provides the first alternative for the algorithm to pursue. Going
from c toe eventually proves useless, and the algorithm has to backtrack from e
to c and then to b. From there, it goes to the vertices f, e, c, and d, from which it
can legitimately return to a, yielding the Hamiltonian circuit a, b, f, e, c, d, a. If
we wanted to find another Hamiltonian circuit, we could continue this process by
backtracking from the leaf of the solution found.
Subset-Sum Problem
As our last example, we consider the subset-sum problem: fmd a subset of a given
setS= {s1, ... , sn) of n positive integers whose sum is equal to a given positive
integer d. For example, for S = {1, 2, 5, 6, 8] and d = 9, there are two solutions: {1,
2, 6] and {1, 8]. Of course, some instances of this problem may have no solutions.
+
420
Coping with the Limitations of Algorithm Power
It is convenient to sort the set's elements in increasing order. So we will assume
that
The state-space tree can be constructed as a binary tree like that in Figure 12.4 for
the instanceS= (3, 5, 6, 7} and d = 15. The root of the tree represents the starting
point, with no decisions about the given elements made as yet. Its left and right
children represent, respectively, inclusion and exclusion of s1 in a set being sought.
Similarly, going to the left from a node of the first level corresponds to inclusion
of s2 , while going to the right corresponds to its exclusion, and so on. Thus, a path
from the root to a node on the ith level of the tree indicates which of the first i
numbers have been included in the subsets represented by that node.
We record the value of s', the sum of these numbers, in the node. If s' is equal
to d, we have a solution to the problem. We can either report this result and stop
or, if all the solutions need to be found, continue by backtracking to the node's
parent. If s' is not equal to d, we can terminate the node as non promising if either
of the following two inequalities holds:
s' + si+! > d (the sums' is too large)
s' +
L" s < d (the sum s' is too small).
j
j=i+l
0
w/o3
with 3
3
with 5
w/o5
0
5
3
8
0
with 5
w/o 5
with 6
w/o6
with 6
w/o6
with 6
w/o 6
14
8
9
3
11
5
X
0+13<15
X
with 7
w/o 7
14+7>15
15
8
solution
X
X
X
X
X
9+7>15
3+7<15
11+7>15
5+7<15
8<15
FIGURE 12.4 Complete state-space tree of the backtracking algorithm applied to the
instance S = {3, 5, 6, 7} and d = 15 of the subset-sum problem. The
number inside a node is the sum of the elements already included in
subsets represented by the node. The inequality below a leaf indicates
the reason for its termination.
12.1
Backtracking
421
General Remarks
From a more general perspective, most backtracking algorithms fit the following description. An output of a backtracking algorithm can be thought of as an
n-tuple (xt, x2 , ... , X11 ) where each coordinate xi is an element of some finite linearly ordered set S;. For example, for then-queens problem, each S; is the set
of integers (column numbers) 1 through n. The tuple may need to satisfy some
additional constraints (e.g., the nona!tacking requirements in then-queens problem). Depending on the problem, all solution tuples can be of the same length
(then-queens and the Hamiltonian circuit problem) or of different lengths (the
subset-sum problem). A backtracking algorithm generates, explicitly or implicitly, a state-space tree; its nodes represent partially constructed tuples with the
first i coordinates defined by the earlier actions of the algorithm. If such a tuple
(x1, x 2 , ... , x;) is not a solution, the algorithm finds the next element in S;+J that
is consistent with the values of (x 1 , x 2 , ... , x 1) and the problem's constraints and
adds it to the tuple as its (i + 1)st coordinate. If such an element does not exist,
the algorithm backtracks to consider the next value of x;, and so on.
To start a backtracking algorithm, the following pseudocode can be called for
i = 0; X[l..O] represents the empty tuple.
AlGORITHM
Backtrack(X[l..i])
//Gives a template of a generic backtracking algorithm
//Input: X[Li] specifies first i promising components of a solution
//Output: All the tuples representing the problem's solutions
if X[Li] is a solution write X[l..i]
else //see Problem 8 in the exercises
for each element x E S;+J consistent with X[Li] and the constraints do
X[i + 1] *-X
Backtrack(X[1..i + 1]J
Our success in solving small instances of three difficult problems earlier in
this section should not lead you to the false conclusion that backtracking is a
very efficient technique. In the worst case, it may have to generate all possible
candidates in an exponentially (or faster) growing state space of the problem at
hand. The hope, of course, is that a backtracking algorithm will be able to prune
enough branches of its state-space tree before running out of time or memory or
both. The success of this strategy is known to vary widely, not only from problem
to problem but also from one instance to another of the same problem.
There are several tricks that might help reduce the size of a state-space tree.
One is to exploit the symmetry often present in combinatorial problems. For
example, the board of then-queens problem has several symmetries so that some
solutions can be obtained from others by reflection or rotation. This implies, in
particular, that we need not consider placements of the first queen in the last l n j2j
columns, because any solution with the first queen in square (1, i), rn/21 co i con,
:
~~~il 1
1
i-'i ill
1.''•.'·'".'.
I"
.'I'
I
422
Coping with the Limitations of Algorithm Power
can be obtained by reflection (which?) from a solution with the first queen in
square (1, n- i + 1). This observation cuts the size of the tree by about half.
Another trick is to preassign values to one or more components of a solution,
as we did in the Hamiltonian circuit example. Data presorting in the subset -sum
example demonstrates potential benefits of yet another opportunity: rearrange
data of an instance given.
It would be highly desirable to be able to estimate the size of the state-space
tree of a backtracking algorithm. As a rule, this is too difficult to do analytically,
however. Knuth [Knu75] suggested generating a random path from the root to
a leaf and using the information about the number of choices available during
the path generation for estimating the size of the tree. Specifically, let c1 be the
number of values of the first component x1 that are consistent with the problem's
constraints. We randomly select one of these values (with equal probability llct)
to move to one of the root's c1 children. Repeating this operation for cz possible
values for x2 that are consistent with x 1 and the other constraints, we move to one
of the c2 children of that node. We continue this process until a leaf is reached
after randomly selecting values for x1, x 2 , ... , x,. By assuming that the nodes on
level i has ci children on average, we estimate the number of nodes in the tree as
Generating several such estimates and computing their average yields a useful
estimation of the actual size of the tree, although the standard deviation of this
random variable can be large.
In conclusion, three things on behalf of backtracking need to be said. First, it
is typically applied to difficult combinatorial problems for which no efficient algorithms for finding exact solutions possibly exist. Second, unlike the exhaustivesearch approach, which is doomed to be extremely slow for all instances of a
problem, backtracking at least holds a hope for solving some instances of nontrivial sizes in an acceptable amount of time. This is especially true for optimization
problems, for which the idea of backtracking can be further enhanced by evaluating the quality of partially constructed solutions. How this can be done is explained
in the next section. Third, even if backtracking does not eliminate any elements
of a problem's state space and ends up generating all its elements, it provides a
specific technique for doing so, which can be of value in its own right.
-----Exercises 12.1 - - - - - - - - - - - - - - 1. a. Continue the backtracking search for a solution to the four-queens problem, which was started in this section, to find the second solution to the
problem.
b. Explain how the board's symmetry can be used to find the second solution
to the four-queens problem.
12.1
Backtracking
423
2. a. Which is the last solution to the five-queens problem found by the backtracking algorithm?
b. Use the board's symmetry to find at least four other solutions to the
problem.
3. a. Implement the backtracking algorithm for the n-queens problem in the
language of your choice. Run your program for a sample of n values to
get the numbers of nodes in the algorithm's state-space trees. Compare
these numbers with the numbers of candidate solutions generated by the
exhaustive-search algorithm for this problem.
b. For each value of n for which you run your program in part (a), estimate
the size of the state-space tree by the method described in Section 12.1 and
compare the estimate with the actual number of nodes you obtained.
4. Apply backtracking to the problem of finding a Hamiltonian circuit in the
following graph.
c
e
5. Apply backtracking to solve the 3-coloring problem for the graph in Figure
12.3a.
6. Generate all permutations of {1, 2, 3, 4] by backtracking.
7. a. Apply backtracking to solve the following instance of the subset-sum problem: S = {1, 3, 4, 5] and d = 11.
b. Will the backtracking algorithm work correctly if we use just one of the
two inequalities to terminate a node as nonpromising?
8. The general template for backtracking algorithms, which was given in Section
11.1, works correctly only if no solution is a prefix to another solution to the
problem. Change the pseudocode to work correctly for such problems as well.
9. Write a program implementing a backtracking algorithm for
a. the Hamiltonian circuit problem.
b. them-coloring problem.
Tills puzzle-like game is played on a triangular board with 15
small holes arranged in an equilateral triangle. In an initial position, all but
one of the holes are occupied by pegs, as in the example shown below. A legal
10. Puzzle pegs
move is a jump of a peg over its immediate neighbor into an empty square
opposite; the jump removes the jumped-over neighbor from the board.
'ii
Coping with the Limitations of Algorithm Power
424
Design and implement a backtracking algorithm for solving the following
versions of this puzzle.
a. Starting with a given location of the empty hole, find a shortest sequence
of moves that eliminates 14 pegs with no limitations on the final position
of the remaining peg.
b. Starting with a given location of the empty hole, find a shortest sequence
of moves that eliminates 14 pegs with the remaining one at the empty hole
of the initial board.
12
Branch-and-Bound
Recall that the central idea of backtracking, discussed in the previous section, is
to cut off a branch of the problem's state-space tree as soon as we can deduce
that it cannot lead to a solution. This idea can be strengthened further if we
deal with an optimization problem, one that seeks to minimize or maximize
an objective function, usually subject to some constraints (a tour's length, the
value of items selected, the cost of an assignment, and the like). Note that in
the standard terminology of optimization problems, a feasible solution is a point
in the problem's search space that satisfies all the problem's constraints (e.g., a
Hamiltonian circuit in the traveling salesman problem or a subset of items whose
total weight does not exceed the knapsack's capacity in the knapsack problem),
while an optimal solution is a feasible solution with the best value of the objective
function (e.g., the shortest Hamiltonian circuit or the most valuable subset ofitems
that fit the knapsack).
Compared to backtracking, branch-and-bound requires two additional items:
"
"
1.
a way to provide, for every node of a state-space tree, a bound on the best
value of the objective function 1 on any solution that can be obtained by adding
further components to the partially constructed solution represented by the
node
the value of the best solution seen so far
This bound should be a lower bound for a minimization problem and an upper bound for a maximization problem.
12.2 Branch-and-Bound
425
If this information is available. we can compare a node's bound value with the
value of the best solution seen so far: if the bound value is not better than the best
solution seen so far-i.e., not smaller for a minimization problem and not larger
for a maximization problem-the node is nonpromising and can be terminated
(some people say the branch is "pruned") because no solution obtained from it
can yield a better solution than the one already available. This is the principal idea
of the branch-and-bound technique.
In general, we terminate a search path at the current node in a state-space
tree of a branch-and-bound algorithm for any one of the following three reasons:
"'
The value of the node's bound is not better than the value of the best solution
seen so far.
"'
The node represents no feasible solutions because the constraints of the
problem are already violated.
"
The subset of feasible solutions represented by the node consists of a single
point (and hence no further choices can be made )-in this case we compare
the value of the objective function for this feasible solution with that of the
best solution seen so far and update the latter with the former if the new
solution is better.
Assignment Problem
Let us illustrate the branch-and-bound approach by applying it to the problem o£
assigning n people to n jobs so that the total cost of the assignment is as small
as possible. We introduced this problem in Section 3.4, where we solved it by
exhaustive search. Recall that an instance of the assignment problem is specified
by an n-by-n cost matrix C so that we can state the problem as follows: select one
element in each row of the matrix so that no two selected elements are in the
same column and their sum is the smallest possible. We will demonstrate how this
problem can be solved using the branch-and-bound technique by considering the
same small instance we investigated in Section 3.4:
job 1
c{
5
7
job2
job 3
job 4
2
4
8
7
3
1
9
l
6
person a
person b
person c
person d
How can we find a lower bound on the cost of an optimal selection without
actually solving the problem? We can do this by several methods. For example, it
is clear that the cost of any solution, including an optimal one, cannot be smaller
than the sum of the smallest elements in each of the matrix's rows. For the instance
here, this sum is 2 + 3 + 1 + 4 = 10. It is important to stress that this is not the cost
of any legitimate selection (3 and 1 came from the same column of the matrix);
it is just a lower bound on the cost of any legitimate selection. We can and will
+
426
Coping with the Limitations of Algorithm Power
apply the same thinking to partially constructed solutions. For example, for any
legitimate selection that selects 9 from the first row, the lower bound will be
9+3+1+4=17.
One more comment is in order before we embark on constructing the problem's state-space tree. It deals with the order in which the tree's nodes will be
generated. Rather than generating a single child of the last promising node as
we did in backtracking, we will generate all the children of the most promising
node among nonterminated leaves in the current tree. (Nonterrninated, i.e., still
promising, leaves are also called live.) How can we tell which of the nodes is most
promising? We can do this by comparing the lower bounds of the live nodes. It
is sensible to consider a node with the best bound as most promising, although
this does not, of course, preclude the possibility that an optimal solution will ultimately belong to a different branch of the state-space tree. This variation of the
strategy is called the best-first branch-and-bound.
Returning to the instance of the assignment problem given earlier, we start
with the root that corresponds to no elements selected from the cost matrix. As
we already discussed, the lower-bound value for the root, denoted lb, is 10. The
nodes on the first level of the tree correspond to selections of an element in the
first row of the matrix, i.e., a job for person a (Figure 12.5).
So we have four live leaves (nodes 1 through 4) that may contain an optimal
solution. The most promising of them is node 2 because it has the smallest lowerbound value. Following our best-first search strategy, we branch out from that
node first by considering the three different ways of selecting an element from the
second row and not in the second column-the three different jobs that can be
assigned to person b (Figure 12.6).
Of the six live leaves (nodes 1, 3, 4, 5, 6, and 7) that may contain an optimal
solution, we again choose the one with the smallest lower bound, node 5. First, we
0
start
3
a -> 3
FIGURE 12.5 Levels 0 and 1 of the state-space tree for the instance of the assignment
problem being solved with the best-first branch-and-bound algorithm. The
number above a node shows the order in which the node was generated.
A node's fields indicate the job number assigned to person a and the
lower bound value, lb, for this node.
12.2
Branch-and-Bound
427
FIGURE 12.6 Levels 0, 1, and 2 of the state-space tree for the instance of the assignment
problem being solved with the best-first branch-and-bound algorithm
consider selecting the third column's element from c's row (i.e., assigning person
c to job 3); this leaves us with no choice but to select the element from the fourth
column of d's row (assigning person d to job 4). This yields leaf 8 (Figure 12.7),
which corresponds to the feasible solution {a-> 2, b-> 1, c-> 3, d-> 4] with the
total cost of 13. Its sibling, node 9, corresponds to the feasible solution {a -> 2,
b-> 1, c-> 4, d-> 3} with the total cost of 25. Since its cost is larger than the cost
8
9
c~3
c---74
d-->4
d--.3
cost=13
cost= 25
solution
inferior solution
FIGURE 12.7 Complete state-space tree for the instance of the assignment problem
solved with the best-first branch-and-bound algorithm
428
Coping with the Limitations of Algorithm Power
of the solution represented by leaf 8, node 9 is simply terminated. (Note that if
its cost were smaller than 13, we would have to replace the information about the
best solution seen so far with the data provided by this node.)
Now, as we inspect each of the live leaves of the last state-space tree (nodes
1, 3, 4, 6, and 7 in Figure 12.7), we discover that their lower-bound values are
not smaller than 13, the value of the best selection seen so far (leaf 8). Hence,
we terminate all of them and recognize the solution represented by leaf 8 as the
optimal solution to the problem.
Before we leave the assignment problem, we have to remind ourselves again
that, unlike for our next examples, there is a polynomial-time algorithm for this
problem called the Hungarian method (e.g., [Pap82]). In the light of this efficient
algorithm, solving the assignment problem by branch-and-bound should be considered a convenient educational device rather than a practical recommendation.
Knapsack Problem
Let us now discuss how we can apply the branch-and-bound technique to solving
the knapsack problem. This problem was introduced in Section 3.4: given 11 items
of known weights w 1 and values v1, i = 1, 2, ... , 11, and a knapsack of capacity W,
find the most valuable subset of the items that fit in the knapsack. It is convenient
to order the items of a given instance in descending order by their value-to-weight
ratios. Then the first item gives the best payoff per weight unit and the last one
gives the worst payoff per weight unit, with ties resolved arbitrarily:
It is natural to structure the state-space tree for this problem as a binary tree
constructed as follows (Figure 12.8). Each node on the ith level of this tree, 0 s
i s 11, represents all the subsets of n items that include a particular selection made
from the first i ordered items. This particular selection is uniquely determined
by the path from the root to the node: a branch going to the left indicates the
inclusion of the next item, while a branch going to the right indicates its exclusion.
We record the total weight w and the total value v of this selection in the node,
along with some upper bound ub on the value of any subset that can be obtained
by adding zero or more items to this selection.
A simple way to compute the upper bound ub is to add to v, the total value of
the items already selected, the product of the remaining capacity of the knapsack
W - w and the best per unit payoff among the remaining items, which is v1+1/ w1+l:
(12.1)
As a specific example, let us apply the branch-and-bound algorithm to the
same instance of the knapsack problem we solved in Section 3.4 by exhaustive
search. (We reorder the items in descending order of their value-to-weight ratios,
though.)
12.2 Branch-and-Bound
item
weight
value
1
2
3
4
7
$40
$42
$25
$12
5
3
4
429
value
weight
10
6
5
4
The knapsack's capacity W is 10.
At the root of the state-space tree (see Figure 12.8), no items have been
selected as yet. Hence, both the total weight o£ the items already selected w and
their total value v are equal to 0. The value of the upper bound computed by
formula (12.1) is $100. Node 1, the left child of the root, represents the subsets
that include item 1. The total weight and value of the items already included are 4
with 1
with 2
w/o 1
w/o2
inferior to
4
W=4, V=40
ub = 70
with 3
node 8
w/o3
5
W= 9, V= 65
ub= 69
inferior to node 8
optimal solution
FIGURE 12.8 State-space tree of the branch-and-bound algorithm for the instance of
the knapsack problem
430
Coping with the Limitations of Algorithm Power
and $40, respectively; the value of the upper bound is 40 + (10- 4) * 6 ~ $76.
Node 2 represents the subsets that do not include item 1. Accordingly, w = 0,
v = $0, and ub ~ 0 + (10- 0) * 6 ~ $60. Since node 1 has a larger upper bound than
the upper bound of node 2, it is more promising for this maximization problem,
and we branch from node 1 first. Its children-nodes 3 and 4-represent subsets
with item 1 and with and without item 2, respectively. Since the total weight w
of every subset represented by node 3 exceeds the knapsack's capacity, node 3
can be terminated immediately. Node 4 has the same values of w and v as its
parent; the upper bound ub is equal to 40 + (10- 4) * 5 ~ $70. Selecting node 4
over node 2 for the next branching (why?), we get nodes 5 and 6 by respectively
including and excluding item 3. The total weights and values as well as the upper
bounds for these nodes are computed in the same way as for the preceding nodes.
Branching from node 5 yields node 7, which represents no feasible solutions, and
node 8, which represents just a single subset (1, 3). (As there are no additional
items to consider, the upper bound for node 7 is simply equal to the total value of
these two items.) The remaining live nodes 2 and 6 have smaller upper-bound
values than the value of the solution represented by node 8. Hence, both can
be terminated making the subset (1, 3) of node 8 the optimal solution to the
problem.
Solving the knapsack problem by a branch-and-bound algorithm has a rather
unusual characteristic. Typically, internal nodes of a state-space tree do not define
a point of the problem's search space, because some of the solution's components
remain undefined. (See, for example, the branch-and-bound tree for the assignment problem discussed in the preceding subsection.) For the knapsack problem,
however, every node of the tree represents a subset of the items given. We can use
this fact to update the information about the best subset seen so far after generating each new node in the tree. If we had done this for the instance investigated
above, we could have terminated nodes 2 and 6 before node 8 was generated,
because they both are inferior to the subset of value $65 of node 5.
Traveling Salesman Problem
We will be able to apply the branch-and-bound technique to instances of the
traveling salesman problem if we come up with a reasonable lower bound on tour
lengths. One very simple lower bound can be obtained by finding the smallest
element in the intercity distance matrix D and multiplying it by the number of
cities n. But there is a less obvious and more informative lower bound, which
does not require a lot of work to compute. It is not difficult to show (Problem 8
in Exercises 12.2) that we can compute a lower bound on the length l of any tour
as follows. For each city i, 1 ::: i ::: n, find the sums; of the distances from city i to
the two nearest cities; compute the sum s of these n numbers; divide the result by
2; and, if all the distances are integers, round up the result to the nearest integer:
lb= [s/21-
(12,2)
~
12.2
Branch-and-Bound
431
I
8
9
3
(a)
8
a, b, c, d,
a, b, d, e,
(e, a)
(c, a)
I= 24
first tour
better tour
I= 16
inferior tour optimal tour
(b)
FIGURE 12.9 (a) Weighted graph. (b) State-space tree of the the branch-and-bound
algorithm to find the shortest Hamiltonian circuit in this graph. The list of
vertices in a node specifies a beginning part of the Hamiltonian circuits
represented by the node.
For example, for the instance in Figure 12.9a, formula (12.2) yields
lb =
rw + 3) + 13 + 6) + 11 + 2) + 13 + 4J + 12 + 3JJ/21 = 14.
Moreover, for any subset of tours that must include particular edges of a given
graph, we can modify lower bound (12.2) accordingly. For example, for all the
Hamiltonian circuits of the graph in Figure 12.9a that must include edge (a, d), we
get the following lower bound by summing the lengths of the two shortest edges
incident with each of the vertices, with the required inclusion of edges (a, d) and
(d, a):
r[(l+ 5) + (3 + 6) + (1+ 2) + (3 + 5) + (2 + 3)]/21 = 16.
We now apply the branch-and-bound algorithm, with the bounding function
given by formula (12.2), to find the shortest Hamiltonian circuit for the graph in
Figure 12.9a. To reduce the amount of potential work, we take advantage of two
observations made in Section 3.4. First, without loss of generality, we can consider
only tours that start at a. Second) because our graph is undirected, we can generate
I
432
Coping with the Limitations of Algorithm Power
only tours in which b is visited before c. In addition, after visiting n - 1 = 4 cities,
a tour has no choice but to visit the remaining unvisited city and return to the
starting one. The state-space tree tracing the algorithm's application is given in
Figure 12.9b.
The comments we made at the end of the preceding section about the strengths
and weaknesses of backtracking are applicable to branch-and-bound as well. To
reiterate the main point: these state-space tree techniques enable us to solve
many large instances of difficult combinatorial problems. As a rule, however, it is
virtually impossible to predict which instances will be solvable in a realistic amount
of time and which will not.
Incorporation of additional information, such as a symmetry of a game's
board, can widen the range of solvable instances. Along this line, a branch-andbound algorithm can be sometimes accelerated by a knowledge of the objective
function's value of some nontrivial feasible solution. The information might be
obtainable-say, by exploiting specifics of the data or even, for some problems,
generated randomly-before we start developing a state-space tree. Then we can
use such a solution immediately as the best one seen so far rather than waiting for
the branch-and-bound processing to lead us to the first feasible solution.
In contrast to backtracking, solving a problem by branch-and-bound has
both the challenge and opportunity of choosing an order of node generation and
finding a good bounding function. Though the best-first rule we used above is a
sensible approach, it may or may not lead to a solution faster than other strategies.
(The branch of computer science called artificial intelligence (AI) is particularly
interested in different strategies for developing state-space trees.)
Finding a good bounding function is usually not a simple task. On the one
hand, we want this function to be easy to compute. On the other hand, it cannot
be too simplistic-otherwise, it would fail in its principal task to prune as many
branches of a state-space tree as soon as possible. Striking a proper balance between these two competing requirements may require intensive experimentation
with a wide variety of instances of the problem in question.
-----Exercises 12
1. What data structure would you use to keep track of live nodes in a best-first
branch-and-bound algorithm?
2. Solve the same instance of the assignment problem as the one solved in
the section by the best-first branch-and-bound algorithm with the bounding
function based on matrix columns rather than rows.
3. a. Give an example of the best-case input for the branch-and-bound algorithm for the assigmnent problem.
~
II
12.2 Branch-and-Bound
433
b. In the best case, how many nodes will be in the state-space tree of the
branch-and-bound algorithm for the assignment problem?
4. Write a program for solving the assignment problem by the branch-and-bound
algorithm. Experiment with your program to determine the average size of
the cost matrices for which the problem is solved in under one minute on
your computer.
5. Solve the following instance of the knapsack problem by the branch-and-
bound algorithm.
item
weight
value
1
2
3
4
10
7
8
4
$100
$63
W=16
$56
$12
6. a. Suggest a more sophisticated bounding function for solving the knapsack
problem than the one used in the section.
b. Use your bounding function in the branch-and-bound algorithm applied
to the instance of Problem 5.
7. Write a program to solve the knapsack problem with the branch-and-bound
algorithm.
8. a. Prove the validity of the lower bound given by formula (12.2) for instances
of the traveling salesman problem with symmetric matrices of integer
intercity distances.
b. How would you modify lower bound (12.2) for nonsymmetric distance
matrices?
9. Apply the branch-and-bound algorithm to solve the traveling salesman problem for the following graph.
2
a )--....C.---{b
5
8
7
3
c)-----{d
(We solved this problem by exhaustive search in Section 3.4.)
434
Coping with the Limitations of Algorithm Power
10. As a research project, write a report on how state-space trees are used for
programming such games as chess, checkers, and tic-tac-toe. The two principal
algorithms you should read about are the minimax algorithm and alpha-beta
pruning.
12
Approximation Algorithms for
NP-hard Problems
In this section, we discuss a different approach to handling difficult problems of
combinatorial optimization, such as the traveling salesman problem and the knapsack problem. As we pointed out in Section 11.3, the decision versions of these
problems are NP-complete. The optimization versions of such difficult combinatorial problems fall in the class of NP-hard problems-problems that are at least
as hard as NP-complete problems2 Hence, there are no known polynomial-time
algorithms for these problems, and there are serious theoretical reasons to believe that such algorithms do not exist. What then are our options for handling
such problems, many of which are of significant practical importance?
If an instance of the problem in question is very small, we might be able to
solve it by an exhaustive-search algorithm (Section 3.4). Some such problems can
be solved by the dynamic programming technique as demonstrated in Section
8.4. But even when this approach works in principle, its practicality is limited by
dependence on the instance parameters being relatively small. The discovery of
the branch-and-bound technique has proved to be an important breakthrough,
because this technique makes it possible to get solutions to many large instances
of difficult problems of combinatorial optimization in an acceptable amount of
time. However, such good performance cannot usually be guaranteed.
There is a radically different way of dealing with difficult optimization problems: solve them approximately by a fast algorithm. This approach is particularly
appealing for applications where a good but not necessarily optimal solution will
suffice. Besides, in real-life applications, we often have to operate with inaccurate
data to begin with. Under such circumstances, going for an approximate solution
can be a particularly sensible choice.
Although approximation algorithms run a gamut in level of sophistication,
most of them are based on some problem-specific heuristic. A heuristic is a
common-sense rule drawn from experience rather than from a mathematically
proved assertion. For example, going to the nearest unvisited city in the traveling
2.
The notion of an NP-hard problem can be defined more formally by extending the notion of polynomial
reducability to problems that arc not necessarily in class NP, including optimization problems of the
type discussed in this section (see [Gar79], Chapter 5).
12.3
Approximation Algorithms for NP-hard Problems
435
salesman problem is a good illustration of this notion. We discuss an algorithm
based on this heuristic later in this section.
Of course, if we use an algorithm whose output is just an approximation of the
actual optimal solution, we would like to know how accurate this approximation
is. We can quantify the accuracy of an approximate solution sa to a problem
minimizing some function f by the size of the relative error of this approximation
re(sa) = f(sa)- f(s*),
f(s*)
where s* is an exact solution to the problem. Alternatively, smce re(sc) =
f(s 0 )/f(s'') -1, we can simply use the accuracy ratio
r(s ) = f(sa)
a
f(s*)
as a measure of accuracy of sa· Note that for the sake of scale uniformity, the
accuracy ratio of approximate solutions to maximization problems is often computed as
r(sa) = f(s')
f(sa)
to make this ratio greater than or equal to 1, as it is for minimization problems.
Obviously, the closer r(s 0 ) is to 1, the better the approximate solution is.
For most instances, however, we cannot compute the accuracy ratio, because we
typically do not know .f(s*), the true optimal value of the objective function.
Therefore, our hope should lie in obtaining a good upper bound on the values
of r(s 0 ). This leads to the following definition.
DEFINITION A polynomial-time approximation algorithm is said to be a capproximation algorithm, where c::: 1, if the accuracy ratio of the approximation
it produces does not exceed c for any instance of the problem in quesiton:
(12.3)
The best (i.e., the smallest) value of c for which inequality (12.3) holds for all
instances of the problem is called the peiformance ratio of the algorithm and
denoted RA·
The performance ratio serves as the principal metric indicating the quality of
the approximation algorithm. We would like to have approximations algorithms
with RA as close to 1 as possible. Unfortunately, as we shall see, some approximation algorithms have infinitely large performance ratios (RA = oo). This does not
necessarily rule out using such algorithms, but it does call for a cautious treatment
of their outputs.
There are two important facts about difficult problems of combinatorial optimization worth keeping in mind. First, although the difficulty level of solving most
.
}, .. I. I
n.l':
I
436
Coping with the Limitations of Algorithm Power
I
such problems exactly is the same to within a polynomial-time transformation of
one problem to another, this equivalence does not translate into the realm of approximation algorithms. Finding approximate solutions with a reasonable level of
accuracy is much easier for some of these problems than for the others. Second,
some of the problems have special classes of instances that are both particularly
important for real-life applications and easier to solve than their general counterparts. The traveling salesman problem is a prime example of this situation.
Approximation Algorithms for the Traveling Salesman
Problem
We solved the traveling salesman problem by exhaustive search in Section 3.4,
mentioned its decision version as one of the most well-known NP-complete problems in Section 11.3, and saw how its instances can be solved by a branch-andbound algorithm in Section 12.2. Here, we consider several approximation algorithms, a small sample of dozens of such algorithms suggested over the years
for this famous problem. (For a much more detailed discussion of the topic, see
[Law85], [Hoc97], [Joh97], and [Gut02].)
But first let us answer the question of whether we should hope to find a
polynomial-time approximation algorithm with a finite performance ratio on all
instances of the traveling salesman problem. As the following theorem [Sah76]
shows, the answer turns out to be no, unless P = N P.
THEOREM 1 If P i" NP, there exists no c-approximation algorithm for the
traveling salesman problem, i.e., there exists no polynomial-time approximation
algorithm for this problem so that for all instances
f(s") :" cf(s')
for some constant c.
PROOF By way of contradiction, suppose that such an approximation algorithm
A and a constant e exist. (Without loss of generality, we can assume that e is a
positive integer.) We will show that this algorithm could then be used for solving
the Hamiltonian circuit problem in polynomial time. We will take advantage of
a variation of the transformation used in Section 11.3 to reduce the Hamiltonian
circuit problem to the traveling salesman problem. Let G be an arbitrary graph
with n vertices. We map G to a complete weighted graph G' by assigning weight
1 to each edge in G and adding an edge of weight en + 1 between each pair of
vertices not adjacent in G. If G has a Hamiltonian circuit, its length in G' is
n; hence, it is the exact solution s' to the traveling salesman problem for G'.
Note that if s" is an approximate solution obtained for G' by algorithm A, then
f(s") :"en by the assumption. If G does not have a Hamiltonian circuitinG, the
shortest tour in G' will contain at least one edge of weight en + 1, and hence
f(sa)?: f(s') >en. Taking into account the two derived inequalities, we could
12.3
Approximation Algorithms for NP.!nard Problems
437
a )----"---{ b
6
2
3
3
dl-----{c
FIGURE 12.10 Instance of the traveling salesman problem for illustrating the nearestneighbor algorithm.
solve the Hamiltonian circuit problem for graph Gin polynomial time by mapping
G toG', applying algorithm A to get tour sa in G'. and comparing its length with en.
Since the Hamiltonian circuit problem is NP-complete, we have a contradiction
unless P = NP.
llil
Nearest-neighbor algorithm
The following simple greedy algorithm is based on the nearest-neighbor heuristic:
the idea of always going to the nearest unvisited city next.
Step 1 Choose an arbitrary city as the start.
Step 2 Repeat the following operation until all the cities have been visited:
go to the unvisited city nearest the one visited last (ties can be broken
arbitrarily).
Step 3 Return to the starting city.
EXAMPLE 1 For the instance represented by the graph in Figure 12.10, with a as
the starting vertex, the nearest -neighbor algorithm yields the tour (Hamiltonian
circuit) sa: a - b - c- d -a of length 10.
The optimal solution, as can be easily checked by exhaustive search, is the tour
s': a - b- d- c -a oflength 8. Thus, the accuracy ratio of this approximation is
r(s ) = /(sa) = 10 = 1.25
a
f(s')
8
(i.e., tour sa is 25% longer than the optimal tours').
•
Unfortunately, except for its simplicity, not many good things can be said
about the nearest-neighbor algorithm. In particular, nothing can be said in general
about the accuracy of solutions obtained by this algorithm because it can force us
to traverse a very long edge on the last leg of the tour. Indeed, if we change the
weight of edge (a, d) from 6 to an arbitrary large number w 2: 6 in Example 1,
the algorithm will still yield the tour a - b- c- d- a of length 4 + w, while the
438
Coping with the Limitations of Algorithm Power
optimal solution will still be a - b - d - c - a of length 8. Hence,
r(s ) = f(sa) = 4 + w,
"
f(s*)
8
which can be made as large as we wish by choosing an appropriately large value
of w. Hence, RA = oo for this algorithm (as it should be according to Theorem 1).
Another natural greedy algorithm for the traveling salesman problem considers it as the problem of finding a minimum-weight collection of edges in a given
complete weighted graph so that all the vertices have degree 2. (With this emphasis on edges rather than vertices, what other greedy algorithm does it remind you
of?) An application of the greedy technique to this problem leads to the following
algorithm [Ben90].
Multifragment-heuristic algorithm
Step 1 Sort the edges in increasing order of their weights. (Ties can be broken
arbitrarily.) Initialize the set of tour edges to be constructed to the
empty set.
Step 2 Repeat this step until a tour of length n is obtained, where n is the
number of cities in the instance being solved: add the next edge on the
sorted edge list to the set of tour edges, provided this addition does not
create a vertex of degree 3 or a cycle of length less than n; otherwise,
skip the edge.
Step 3 Return the set of tour edges.
As an example, applying the algorithm to the graph in Figure 12.10 yields
[(a, b), (c, d), (b, c), (a, d)}. Thissetofedgesformsthesametourastheoneproduced by the nearest-neighbor algorithm. In general, the multifragment-heuristic
algorithm tends to produce significantly better tours than the nearest-neighbor
algorithm, as we are going to see from the experimental data quoted at the end of
this section. But the performance ratio of the multifragment-heuristic algorithm
is also unbounded, of course.
There is, however, a very important subset of instances, called Euclidean, for
which we can make a nontrivial assertion about the accuracy of both the nearestneighbor and mutifragment-heuristic algorithms. These are the instances in which
intercity distances satisfy the following natural conditions:
"
11
triangle inequality
d[i, j] :s d[i, k] + d[k, j]
for any triple of cities i, j, and k
(the distance between cities i and j cannot exceed the length of a two-leg path
from i to some intermediate city k to j);
symmetry
d[i, j] = d[j, i] for any pair of cities i and j
(the distance from ito j is the same as the distance from j to i).
12.3 Approximation Algorithms for NP-hard Problems
439
A substantial majority of practical applications of the traveling salesman problem are its Euclidean instances. They include, in particular, geometric ones, where
cities correspond to points in the plane and distances are computed by the standard
Euclidean formula. Although the performance ratios of the nearest-neighbor and
multifragment-heuristics algorithms remain unbounded on Euclidean instances,
their accuracy ratios satisfy the following inequality for any such instance with
n ~ 2 cities:
f(sa)
1
f(s') - 2
- - < -(flog2 nl + 1),
where f(sa) and f(s*) are the lengths of the heuristic tour and shortest tour,
respectively (see [Ros77] and [Ong84]).
Minimum-spanning-tree-based algorithms There are approximation algorithms for the traveling salesman problem that exploit a connection between
Hamiltonian circuits and spanning trees of the same graph. Since removing an
edge from a Hamiltonian circuit yields a spanning tree, we can expect that the
structure of a minimum spanning tree provides a good basis for constructing a
shortest tour approximation. Here is an algorithm that implements this idea in a
rather straightforward fashion.
Twice-around-the-tree algorithm
Step 1 Construct a minimum spanning tree of the graph corresponding to a
given instance of the traveling salesman problem.
Step 2 Starting at an arbitrary vertex, perform a walk around the minimum
spanning tree recording all the vertices passed by. (This can be done
by a DFS traversal.)
Step 3 Scan the vertex list obtained in Step 2 and eliminate from it all repeated
occurrences of the same vertex except the starting one at the end of
the list. (This step is equivalent to making shortcuts in the walk.) The
vertices remaining on the list will form a Hamiltonian circuit, which is
the output of the algorithm.
EXAMPLE 2 Let us apply this algorithm to the graph in Figure 12.1la. The
minimum spanning tree of this graph is made up of edges (a, b), (b, c), (b, d),
and (d, e) (Fig. 12.1lb). A twice-around-the-tree walk that starts and ends at a is
a, b, c, b, d,
e, d, b, a.
Eliminating the second b (a shortcut from c to d), the second d, and the third b (a
shortcut from e to a) then yields the Hamiltonian circuit
a, b, c, d, e, a
of length 39.
--,
!
440
Coping with the Limitations of Algorithm Power
a -----------
8
e
11
6
10
c
(b)
(a)
FIGURE 12.11 Illustration of the twice-around-the-tree algorithm. (a) Graph. (b) Walk
around the minimum spanning tree with the shortcuts.
The tour obtained in Example 2 is not optimal. Although that instance is small
enough to find an optimal solution by either exhaustive search or branch-andbound, we refrained from doing so to reiterate a general point. As a rule, we do
uot know what the length of an optimal tour actually is, and therefore we cannot
compute the accuracy ratio .f(sa)/f(s'). For the twice-around-the-tree algorithm,
we can at least estimate it above, provided the graph is Euclidean.
The twice-around-the-tree algorithm is a 2-approximation algorithm for the traveling salesman problem with Euclidean distances.
THEOREM 2
PROOF Obviously, the twice-around-the-tree algorithm is polynomial time if we
use a reasonable algorithm such as Prim's or Kruskal's in Step 1. We need to show
that for any Euclidean instance of the traveling salesman problem, the length of a
tour sa obtained by the twice-around-the-tree algorithm is at most twice the length
of the optimal tour s*; that is,
f(sa) :":: 2f(s*).
Since removing any edge from s* yields a spanning tree T of weight w(T), which
must be greater than or equal to the weight of the graph's minimum spanning tree
w(T*), we get the inequality
.f(s*) > w(T) ": w(T').
This inequality implies that
2.f(s') > 2w(T*) = the length of the walk obtained in Step 2 of the algorithm.
12.3
Approximation Algorithms for NP-hard Problems
441
The possible shortcuts outlined in Step 3 of the algorithm to obtaiu sa cannot
iucrease the total length of the walk in a Euclidean graph; that is,
the length of the walk obtained in Step 2 2: the length of the tour sa.
Combining the last two iuequalities, we get the inequality
2f(s'') > f(sa),
which is, in fact, a slightly stronger assertion than the one we needed to prove.
11!1
Christofides algorithm There is an approximation algorithm with a better performance ratio for the Euclidean traveling salesman problem-the well-known
Christofides algorithm [Chr76]. It also uses a minimum spanning tree but does
this in a more sophisticated way than the twice-around-the-tree algorithm. Note
that a twice-around-the-tree walk generated by the latter algorithm is an Eulerian circuit in the multigraph obtaiued by doubling every edge in the graph given.
Recall that an Eulerian circuit exists in a connected multigraph if and only if all
its vertices have even degrees. The Christofides algorithm obtains such a multigraph by addiug to the graph the edges of a minimum-weight matchiug of all the
odd-degree vertices in its minimum spanning tree. (The number of such vertices
is always even and hence this can always be done.) Then the algorithm finds an
Eulerian circuit in the multigraph and transforms it into a Hamiltonian circuit by
shortcuts, exactly the same way it is done in the last step of the twice-around-thetree algorithm.
EXAMPLE 3 Let us trace the Christofides algorithm in Figure 12.12 on the same
iustance (Figure 12.12a) used for tracing the.twice-around-the-tree algorithm in
Figure 12.11. The graph's minimum spanning tree is shown iu Figure 12.12b. It has
four odd-degree vertices: a, b, c, and e. The minimum-weight matching of these
four vertices consists of edges (a, b) and (c, e). (For this tiny iustance, it can be
found easily by compariug the total weights of just three alternatives: (a, b) and
(c, e), (a, c) and (b, e), (a, e) and (b, c).) The traversal of the multigraph, startiug
at vertex a, produces the Eulerian circuit a - b - c - e - d - b - a, which, after
one shortcut, yields the tour a - b - c - e - d- a of length 37.
T11e performance ratio of the Christofides algorithm on Euclidean iustances
is 1.5 (see, e.g., [Pap82]). It tends to produce significantly better approximations
to optimal tours than the twice-around-the-tree algorithm does iu empirical tests.
(We quote some results of such tests at the end of this subsection.) The quality of
a tour obtaiued by this heuristic can be further improved by optimizing shortcuts
made on the last step of the algorithm as follows: examine the multiply-visited
cities iu some arbitrary order and for each make the best possible shortcut. This
enhancement would not have improved the tour a - b- c- e- d- a obtained in
442
Coping with the Limitations of Algorithm Power
8
11
6
10
c
(a)
I
I
I
I
4 I
I
I
I
I
b
8
6
(b)
''
' \11
d
------- '
/
c
a
I
7
4
''
-' '
e
a
'
I
I
I
I
I
I
I
I
e
4
7
9
11
d
b
6
c
(c)
FIGURE 12.12 Application of the Christofides algorithm. (a) Graph. (b) Minimum
spanning tree with added edges (in dash) of a minimum-weight matching
of all odd-degree vertices. (c) Hamiltonian circuit obtained.
Example 3 from a - b - c - e - d - b - a because shortcutting the second occurrence of b happens to be better than shortcutting its first occurrence. In general,
however, this enhancement tends to decrease the gap between the heuristic and
optimal tour lengths from about 15% to about 10%, at least for randomly generated Euclidean instances [Joh02].
Local search heuristics For Euclidean instances, surprisingly good approximations to optimal tours can be obtained by iterative-improvement algorithms, which
are also called local search heuristics. The best-known of these are the 2-opt, 3opt, and Lin-Kernighan algorithms. These algorithms start with some initial tour,
e.g., constructed randomly or by some simpler approximation algorithm such as
the nearest-neighbor. On each iteration, the algorithm explores a neighborhood
of the current tour by replacing a few edges in the current tour by other edges.
If the changes produce a shorter tour, the algorithm makes it the current tour
and continues by exploring its neighborhood in the same manner; otherwise, the
current tour is returned as the algorithm's output and the algorithm stops.
i
~
12.3
Approximation Algorithms for NP-hard Problems
c,
c,
c,
443
c,
c,
{a)
c,
(b)
FIGURE 12.13 2-change: (a) Original tour. (b) New tour.
The 2-opt algorithm works by deleting a pair of nonadjacent edges in a tour
and reconnecting their endpoints by the different pair of edges to obtain another
tour (see Figure 12.13). This operation is called the2-change. Note that there is
only one way to reconnect the endpoints because the alternative produces two
disjoint fragments.
If we start with the nearest-neighbor tour a - b- c- d - e -a in
the graph in Figure 12.11, whose length l,, is equal to 39, the 2-opt algorithm will
I!
move to the next tour as shown in Figure 12.14.
EXAMPLE 4
To generalize the notion of the 2-change, one can consider the k-change for
any k :0: 2. This operation replaces up to k edges in a current tour. In addition to
2-changes, only the 3-changes have proved to be of practical interest. The two
principal possibilities of 3-changes are shown in Figure 12.15.
There are several other local search algorithms for the traveling salesman
problem. The most prominent of them is the Lin-Kernighan algorithm [Liu73],
which for two decades after its publication in 1973 was considered the best algorithm to obtain high-quality approximations of optimal tours. The Lin-Kernighan
algorithm is a variable-opt algorithm: its move can be viewed as a 3-opt move
followed by a sequence of 2-opt moves. Because of its complexity, we have to refrain from discussing this algorithm here. The excellent surveys by Johnson and
McGeoch ([Joh97], [Joh02]) contain an outline of the algorithm and its modern
extensions, as well as methods for its efficient implementation. These surveys
also contain results from the important empirical studies about performance of
many heuristics for the traveling salesman problem, including, of course, the LinKernighan algorithm. We conclude our discussion by quoting some of these data.
Empirical results The traveling salesman problem has been the subject of intense study for the last fifty years. This interest was driven by a combination of
pure theoretical interest and serious practical needs stemming from such newer
444
Coping with the Limitations of Algorithm Power
12
a
12
a
e
e
I
I
41I
7
7
l = 42 > Inn = 39
I
I
~-
d
//1o
6
c
~e
I
I
I
I
41I
:7
I
I
I
I
v
9
l = 46 > Inn = 39
d
b
6
0
10
c
12
a
9
12
a
e
e
I
I
4
:7
l = 45 > Inn = 39
11
4
I
I
b
a
·~)
12
------
e
a
7
4
b
e
8
9
7
I = 38 < lnn = 39
(new tour)
d
b
'
6',
10
d
10
c
FIGURE 12.14 2-changes from the nearest-neighbor tour of the graph in Figure 12.11
I
I
I:
I
12.3
Approximation Algorithms for NP-hard Problems
445
c,
c,
c,
(a)
c,
c,
(b)
c,
c,
c,
(c)
FIGURE 12.15 3-change: (al Original tour. (b). (cl New tours.
applications as circuit-board and VLSI-chip fabrication. X-ray crystallography,
and genetic engineering. Progress in developing effective heuristics, their efficient
implementation by using sophisticated data structures, and the ever-increasing
power of computers have led to a situation that differs drastically from a pessimistic picture painted by the worst-case theoretical results. This is especially
true for the most important applications class of instances of the traveling salesman problem: points in the two-dimensional plane with the standard Euclidean
distances between them.
~
1)1
446
Coping with the Limitations of Algorithm Power
I
I
It used to be the case that instances with a few hundred cities were considered
too large to be solved exactly. Nowadays, instances with up to 1,000 cities can be
solved exactly in a quite reasonable amount of time-typically, in minutes or less
on a good workstation-by such optimization packages as Concord [App]. In fact,
according to the information on the Web site maintained by the authors of that
package, the largest instance of the traveling salesman problem solved exactly
as of May 2004 was the shortest tour through all 24,978 cities in Sweden. There
should be little doubt that this record will eventually be superseded and our ability
to solve ever larger instances exactly will continue to expand. This remarkable
progress does not eliminate the usefulness of approximation algorithms for such
problems, however. First, some applications lead to instances that are still too large
to be solved exactly in a reasonable amount of time. Second, one may well prefer
spending seconds to find a tour that is within a few percent of optimum than to
spend many hours or even days of computing time to find the shortest tour exactly.
But how can one tell how good or bad the approximate solution is if we do not
know the length of an optimal tour? A convenient way to overcome this difficulty
is to solve the linear programming problem describing the instance in question by
ignoring the integrality constraints. This provides a lower bound-called theHeldKwp bound-on the length of the shortest tour. The Held-Karp bound is typically
very close (less than J%) to the length of an optimal tour, and this bound can be
computed in seconds or minutes unless the instance is truly huge. Thus, for a tour
s, obtained by some heuristic, we estimate the accuracy ratio r(s,) = f(s,)/f(s')
from above by the ratio f(s,)/ H K (s*), where f(s,) is the length of the heuristic
tours, and H K (s') is the Held-Karp lower bound on the shortest-tour length.
The results (see Table 12.1) from a large empirical study [Joh02] indicate the
average tour quality and running times for the discussed heuristics3 . The instances
in the reported sample have 10,000 cities generated randomly and uniformly as
integral-coordinate points in the plane, with the Euclidean distances rounded
to the nearest integer. The quality of tours generated by the heuristics remain
about the same for much larger instances (up to a million cities) as long as they
belong to the same type of instances. The running times quoted are for expert
implementations run on a Compaq ES40 with 500 Mhz Alpha processors and 2
gigabytes of main memory or its equivalents.
Approximation Algorithms for the Knapsack Problem
The knapsack problem, another well-known NP-hard problem, was also introduced in Section 3.4: given n items of known weights w1, ... , W11 and values
v1, ... , v, and a knapsack of weight capacity W, find the most valuable sub3.
We did not include the results for the twice-around-the-tree heuristic because of the inferior quality
of its approximations with the average excess of about 40%. Nor did we quote the results for the
most sophisticated local search heuristics with the average excess over optimum of less than a fraction
of 1%.
I
»'
T
.
!
I
12.3
Approximation Algorithms for NP-hard Problems
447
TABLE 12.1 Average tour quality and running times for
various heuristics on the 10,000-city random
uniform Euclidean instances [Joh02]
0
Heuristic
nearest neighbor
multifragment
Christofides
2-opt
3-opt
Lin-Kernighan
/o excess over the
Held-Karp bound
Running time
(seconds)
24.79
16.42
9.81
0.28
0.20
1.04
1.41
1.50
2.06
4.70
2.88
2.00
set of the items that fits into the knapsack. We saw how this problem can be
solved by exhaustive search (Section 3.4), dynamic programming (Section 8.4),
and branch-and-bound (Section 12.2). Now we will solve this problem by approximation algorithms.
Greedy algorithms for the knapsack problem We can think of several greedy
approaches to this problem. One is to select the items in decreasing order of
their weights; however, heavier items may not be the most valuable in the set.
Alternatively, if we pick up the items in decreasing order of their value, there is
no guarantee that the knapsack's capacity will be used efficiently. Can we find a
greedy strategy that takes into account both the weights and values? Yes, we can,
by computing the value-to-weight ratios v,jw,, i = 1, 2, ... , n, and selecting the
items in decreasing order of these ratios. (In fact, we already used this approach in
designing the branch-and-bound algorithm for the problem in Section 12.2.) Here
is the algorithm based on this greedy heuristic.
Greedy algorithm for the discrete knapsack problem
Step 1 Compute the value-to-weight ratios r; = v,jw,, i = 1, ... , n, for the
items given.
Step 2 Sort the items in nonincreasing order of the ratios computed in Step
1. (Ties can be broken arbitrarily.)
Step 3 Repeat the following operation until no item is left in the sorted list:
if the current item on the list fits into the knapsack, place it in the
knapsack; otherwise, proceed to the next item.
EXAMPLE 5 Let us consider the instance of the knapsack problem with the
knapsack's capacity equal to 10 and the item information as follows:
448
Coping with the Limitations of Algorithm Power
item
weight
value
1
2
3
4
7
3
4
$42
$12
$40
$25
5
Computing the value-to-weight ratios and sorting the items in nonincreasing order
of these efficiency ratios yields
item
weight
value
1
4
7
5
3
$40
$42
$25
$12
2
3
4
value
weight
10
6
5
4
The greedy algorithm will select the first item of weight 4, skip the next item of
weight 7, select the next item of weight 5, and skip the last item of weight 3. The
solution obtained happens to be optimal for this instance (see Section 12.2, where
we solved the same instance by the branch-and-bound algorithm).
1111
Does this greedy algorithm always yield an optimal solution? The answer, of
course, is no: if it did, we would have a polynomial-time algorithm for the NPhard problem. In fact, the following example shows that no finite upper bound on
the accuracy of its approximate solutions can be given either.
EXAMPLE 6
item
weight
value
1
2
1
2
w
w
value
weight
2
1
The knapsack's capacity is W > 2.
Since the items are already ordered as required, the algorithm takes the first
item and skips the second one; the value of this subset is 2. The optimal selection
is item 2, whose value is W. Hence, the accuracy ratio r(sa) of this approximate
solution is W /2, which is unbounded above.
1111
12.3
Approximation Algorithms for NP-hard Problems
449
It is surprisingly easy to tweak this greedy algorithm to get an approximation
algorithm with a finite performance ratio. All it takes is to choose the better of
two alternatives: the one obtained by the greedy algorithm or the one consisting
of a single item of the largest value that fits into the knapsack. (Note that for
the instance of the preceding example, the second alternative is better than the
first one.) It is not difficult to prove that the performance ratio of this enhanced
greedy algorithm is 2. That is, the value of an optimal subsets* will never be more
than twice as large as the value of the subset sa obtained by this enhanced greedy
algorithm, and 2 is the smallest multiple for which such an assertion can be made.
It is instructive to consider the continuous version of the knapsack problem,
as well. In this version, we are permitted to take arbitrary fractions of the items
given. For this version of the problem, it is natural to modify the greedy algorithm
as follows.
Greedy algorithm for the continuous knapsack problem
Step 1 Compute the value-to-weight ratios v;!w,, i = 1, ... , n. for the items
given.
Step 2 Sort the items in nonincreasing order of the ratios computed in Step
1. (Ties can be broken arbitrarily.)
Step 3 Repeat the following operation until the knapsack is filled to its full
capacity or no item is left in the sorted list: if the current item on the
list fits into the knapsack in its entirety, take it and proceed to the next
item; otherwise, take its largest fraction to fill the knapsack to its full
capacity and stop.
For example, for the four-item instance used in Example 5 to illustrate the
greedy algorithm for the discrete version, the algorithm will take the first item of
weight 4 and then 6/7 of the next item on the sorted list to fill the knapsack to its
full capacity.
It should come as no surprise that this algorithm always yields an optimal
solution to the continuous knapsack problem. Indeed, the items are ordered
according to their efficiency in using the knapsack's capacity. If the first item on
the sorted list has weight w1 and value v1, no solution can use w1 units of capacity
with a higher payoff than v1 . If we cannot fill the knapsack with the first item
or its fraction, we should continue by taking as much as we can of the secondmost efficient item, and so on. A formal rendering of this proof idea is somewhat
involved, and we will leave it for the exercises.
Note also that the optimal value of the solution to an instance of the continuous knapsack problem can serve as an upper bound on the optimal value of the
discrete version of the same instance. This observation provides a more sophisticated way of computing upper bounds for solving the discrete knapsack problem
by the branch-and-bound method than the one used in Section 12.2.
450
Coping with the Limitations of Algorithm Power
Approximation schemes We now retnrn to the discrete version of the knapsack
problem. For this problem, unlike the traveling salesman problem, there exist
polynomial-time approximation schemes, which are parametric families of algorithms that allow us to get approximations s~k) with any predefined accuracy level:
f
(s*) ::=:
~
f(sa )
1 + 1/ k f or any mstance
·
· n,
o f size
where k is an integer parameter in the range 0::: k < n. The first approximation
scheme was suggested by S. Sahni in 1975 [Sah75]. This algorithm generates all
subsets of k items or less, and for each one that fits into the knapsack, it adds
the remaining items as the greedy algorithm would (i.e., in nonincreasing order
of their value-to-weight ratios). The subset of the highest value obtained in this
fashion is returned as the algorithm's output.
EXAMPLE 7 A small example of an approximation scheme with k = 2 is provided in Figure 12.16. The algorithm yields (1, 3, 4}, which is the optimal solution
for this instance.
1111
You can be excused for not being overly impressed by this example. And,
indeed, the importance of this scheme is mostly theoretical rather than practical.
It lies in the fact that, in addition to approximating the optimal solution with any
predefined accuracy level, the time efficiency of this algorithm is polynomial inn.
subset added items value
0
item weight value
1
2
3
4
4
7
5
I
$40
$42
$25
$ 4
capacity W = 10
value
{1}
weight
[2}
[3}
{4}
{1, 2}
{1, 3}
{1, 4}
[2, 3}
{2, 4}
{3, 4}
10
6
5
4
(a)
1,3,4
3,4
4
1, 4
1, 3
not feasible
4
3
not feasible
1
$69
$69
$46
$69
$69
$69
$69
$46
$69
(b)
FIGURE 12.16 Example of applying Sahni"s approximation scheme fork= 2. (a} Instance.
(b) Subsets generated by the algorithm.
r
\
12.3 Approximation Algorithms for NP-hard Problems
451
Indeed, the total number of subsets the algorithm generates before adding extra
elements is
k
k
.
k
k
jdl
J~O
L (~) = L n(n -1) .. ·.,(n- J + 1) :<:: L nl :<:: L nk = (k + 1)1l.
J~O
}
J~O
}
For each of those subsets, it needs 0 (n) time to determine the subset's possible
extension. Thus, the algorithm's efficiency is in O(knk+ 1). Note that while being
polynomial inn, the time efficiency of Sahni's scheme is exponential ink. More
sophisticated approximation schemes, called fully polynomial schemes, do not
have this shortcoming. Among several books that discuss such algorithms, the
monographs [Mar90] amd [Kel04] are especially recommended for their wealth
of other material about the knapsack problem.
-----Exercises 1 2 . 3 - - - - - - - - - - - - - - 1. a. Apply the nearest-neighbor algorithm to the instance defined by the dis·
tance matrix below. Start the algorithm at the first city, assuming that the
cities are numbered from 1 to 5.
0 14
14 0
4 5
[ 10 8
00
7
~ ¥
zl
9 0 32
16 32 0
b. Compute the accuracy ratio of this approximate solution.
2. a. Write a pseudocode for the nearest-neighbor algorithm. Assume that its
input is given by an n-by-n distance matrix.
b. What is the time efficiency of the nearest-neighbor algorithm?
3. Apply the twice-around-the-tree algorithm to the graph in Figure 12.11a with
a walk around the minimum spanning tree that starts at the same vertex a but
differs from the walk in Figure 12.1lb. Is the length of the obtained tour the
same as the length of the tour in Figure 12.11b?
4. Prove that making a shortcut of the kind used by the twice-around-the-tree
algorithm cannot increase the tour's length in an Euclidean graph.
5. What is the time efficiency class of the greedy algorithm for the knapsack
problem?
6. Prove that the performance ratio RA of the enhanced greedy algorithm for
the knapsack problem is equal to 2.
7. Consider the greedy algorithm for the bin-packing problem, which is called
the first-fit (FF) algorithm: place each of the items in the order given into the
+
i
452
Coping with the Limitations of Algorithm Power
first bin the item fits in; when there are no such bins, place the item in a new
bin and add this bin to the end of the bin list
a. Apply FF to the instance
s1 = 0.4, s2 = 0.7, s 3 = 0.2, s4 = 0.1, ss = 0.5
and determine whether the solution obtained is optimal.
b. Determine the worst-case time efficiency of FF.
c. Prove that FF is a 2-approximation algorithm.
8. The first-fit decreasing (FFD) approximation algorithm for the bin-packing
problem starts by sorting the items in nonincreasing order of their sizes and
then acts as the first-fit algorithm.
a. Apply FFD to the instance
s 1 = 0.4, s2 = 0.7, s3 = 0.2, s4 = 0.1, s5 = 0.5
and determine whether the solution obtained is optimal.
b. Does FFD always yield an optimal solution? Justify your answer.
c. Prove that FFD is a 1.5-approximation algorithm.
d. Run an experiment to determine which of the two algorithms-FF or
FFD-yields more accurate approximations on a random sample of the
problem's instances.
9. a. Design a simple 2-approximation algorithm for finding a minimum vertex
cover (a vertex cover with the smallest number of vertices) in a given graph.
b. Consider the following approximation algorithm for finding a maximum
independent set (an independent set with the largest number of vertices) in
a given graph. Apply the 2-approximation algorithm of part a and output
all the vertices that are not in the obtained vertex cover. Can we claim that
this algorithm is a 2-approximation algorithm, too?
10. a. Design a polynomial-time greedy algorithm for the graph-coloring problem.
b. Show that the performance ratio of your approximation algorithm is infinitely large.
12.4 Algorithms for Solving Nonlinear Equations
In this section, we discuss several algorithms for solving nonlinear equations in
one unknown,
f(x) = 0.
(12.4)
12.4
Algorithms for Solving Nonlinear Equations
453
There are several reasons for this choice among subareas of numerical analysis.
First of all, this is an extremely important problem from the practical and theoretical points of view. It arises as a mathematical model of numerous phenomena
in the sciences and engineering, both directly and indirectly. (Recall, for example,
that the standard calculus technique for finding extremum points of a function
f(x) is based on finding its critical points, which are the roots of the equation
J'(x) = 0.) Second, it represents the most accessible topic in numerical analysis
and, at the same time, exhibits its typical tools and concerns. Third, some methods for solving equations closely parallel algorithms for array searching and hence
provide examples of applying general algorithm design techniques to problems of
continuous mathematics.
Let us start with dispelling a misconception you might have about solving
equations. Your experience with equation solving from middle school to calculus
courses might have led you to believe that we can solve equations by "factoring"
or by applying a readily available formula. Sorry to break it to you, but you have
been deceived (with the best of educational intentions, of course): you were able
to solve all those equations only because they had been carefully selected to make
it possible. In general, we cannot solve equations exactly and need approximatiou
algorithms to do so.
This is true even for solving the quadratic equation
ax
2
+ hx + c = 0
because the standard formula for its roots
2a
requires computing the square root, which can be done only approximately for
most positive numbers. In addition, as we discussed in Section 11.4, this canonical
formula needs to be modified to avoid the possibility of low-accuracy answers.
What about formulas for roots of polynomials of degrees higher than two?
Such formulas for third- and fourth-degree polynomials exist, but they are too
cumbersome to be of practical value. For polynomials of degrees higher than
four, there can be no general formula for their roots that would involve only the
polynomial's coefficients, arithmetical operations, and radicals (taking roots). This
remarkable result was published first by the Italian mathematician and physician
Paolo Ruffini (1765-1822) in 1799 and rediscovered a quarter century later by the
Norwegian mathematician Niels Abel (1802-1829); it was developed further by
the French mathematician Evariste Galois (1811-1832). 4
4.
Ruffini's discovery was completely ignored by almost all prominent mathematicians of that time. Abel
died young after a difficult life of poverty. Galois was killed in a duel when he was only 21 years old.
Their results on the solution of higher-degree equations are now considered to be among the crowning
achievements in the history of mathematics.
454
Coping with the Limitations of Algorithm Power
The impossibility of such a formula can hardly be considered a great disappointment. As the great German mathematician Carl Friedrich Gauss (1777-1855)
put it in his thesis of 1801, the algebraic solution of an equation was no better than
devising a symbol for the root of the equation and then saying that the equation
had a root equal to the symbol (0Co98].
We can interpret solutions to equation (12.4) as points at which the graph
of the function .f(x) intersects with the x-axis. The three algorithms we discuss
in this section take advantage of this interpretation. Of course, the graph of .f (x)
may intersect the x-axis at a single point (e.g., x 3 = 0), at multiple or even infinitely
many points (sin x = 0), or at no point (ex+ 1 = 0). Equation (12.3) would then
have a single root, several roots, and no roots, respectively. It is a good idea to
sketch a graph of the function before starting to approximate its roots. It can help
to determine the number of roots and their approximate locations. In general, it
is a good idea to isolate roots, i.e., to identify intervals containing a single root of
the equation in question.
Bisection Method
This algorithm is based on an observation that the graph of a continuous function
must intersect with the x-axis between two points a and b at least once if the
function's values have opposite signs at these two points (Figure 12.17).
The validity of this observation is proved as a theorem in calculus courses, and
we take it for granted here. It serves as the basis of the following algorithm, called
the bisection method, for solving equation (12.4). Starting with an interval [a, b]
at whose endpoints f(x) has opposite signs, the algorithm computes the value of
.f(x) at the middle point x,;d = (a+ h)/2. If .f (x,;d) = 0, a root was found and the
algorithm stops. Otherwise, it continues the search for a root either on [a, xmid] or
on [xmid• b ], depending on which of the two halves the values of .f(x) have opposite
signs at the endpoints of the new interval.
Since we cannot expect the bisection algorithm to stumble on the exact value
of the equation's root and stop, we need a different criterion for stopping the algorithm. We can stop the algorithm after the interval [an, h,] bracketing some rootx'
becomes so small that we can guarantee that the absolute error of approximating
fix)
--~-----,L-----~----------~b~--.x
a1
I
I
x1
FIGURE 12.17 First iteration of the bisection method: x 1 is the middle point of interval
[a, b].
12.4 Algorithms for Solving Nonlinear Equations
455
x' by x,,. the middle point of this interval, is smaller than some small preselected
number e > 0. Since x, is the middle point of [a,, b,] and x* lies within this interval
as well, we have
IX - x*l < bn -an,
2
II
-
(U.S)
Hence, we can stop the algorithm as soon as (b,- a,)/2 < e or, equivalently,
(12.6)
It is not difficult to prove that
*
bl - a1
lx, - x I :<: ~ for n = 1, 2, ....
(12.7)
This inequality implies that the sequence of approximations {x,) can be made as
close to root x* as we wish by choosing n large enough. In other words, we can say
that {x,) converges to root x*. Note, however, that because any digital computer
represents extremely small values by zero (Section 11.4), the convergence assertion is true in theory but not necessarily in practice. In fact, if we choose£ below
a certain machine-dependent threshold, the algorithm may never stop! Another
source of potential complications is round-off errors in evaluating values of the
function in question. Therefore it is a good practice to include in a program implementing the bisection method a limit on the number of iterations the algorithm is
allowed to run.
Here is a pseudocode for the bisection method.
ALGORITHM
Bisection(f(x), a, h, eps, N)
//Implements the bisection method for finding a root of f(x) = 0
//Input: Two real numbers a and b, a< h,
II
a continuous function f(x) on [a, b], f(a)f(b) < 0,
II
an upper bound on the absolute error eps > 0,
II
an upper bound on the number of iterations N
//Output: An approximate (or exact) value x of a root in (a, b)
//or an interval bracketing tbe root if the iteration number limit is reached
n +-- 1 //iteration count
while n :<: N do
x +--(a+ b)/2
ifx- a< eps returnx
fval +-- f(x)
if.fval = 0 return x
if.fval * f(a) < 0
b+--x
else a +- x
n+--n+1
return "iteration limit", a, b
456
Coping with the Limitations of Algorithm Power
y
FIGURE 12.18 Graph of function f(x) = x 3 - x -1
Note that we can use inequality (12.7} to find in advance the number of iterations that should suffice, at least in theory, to achieve a preselected accuracy level.
Indeed, choosing the number of iterations 11 large enough to satisfy inequality
(bt- at)/2" < e, i.e.,
b 1 -at
11 > log 2 - - £
(12.8)
does the trick.
EXAMPLE 1 Let us consider equation
x 3 - x -1 = 0.
(12.9)
It has one real root. (See Figure 12.18 for the graph of .f(x) = x 3 - x- 1.) Since
.f(O) < 0 and .f(2) > 0, the root must lie within interval (0, 2). If we choose the
error tolerance level as E = 10-2 , inequality (12.8) would require n > log2 (2/1 o-2 )
or n :::: 8 iterations.
Figure 12.19 contains a trace of the first eight iterations of the bisection
algorithm applied to solving equation (12.9). Thus, we obtained x8 = 1.3203125
as an approximate value for the root x' of equation (12.9), and we can guarantee
that
11.3203125- x'l < w-2
Moreover, if we take into account the signs of the left -hand side of equation (12.9}
at a 8 , b8 , and x8 , we can assert that the root lies between 1.3203125 and 1.328125.
II
The principal weakness of the bisection method as a general algorithm for
solving equations is its slow rate of convergence compared with other known
methods. It is for this reason that the method is rarely used. Also, it cannot be
12.4
457
Algorithms for Solving Nonlinear Equations
b,
n
an
1
2
3
4
5
6
7
8
0.01.0-
2.0+
2.0+
UJ-
1.5+
1.251.251.31251.31251.3125-
1.5+
1.375+
1.375+
1.34375+
1.328125+
x,
f(x,)
1.0
1.5
1.25
1.375
1.3125
1.34375
1.328125
1.3203125
-1.0
0.875
-0.296875
0.224609
-0.051514
0.082611
0.014576
-0.018711
FIGURE 12.19 Trace of the bisection method for solving equation (12.91. The signs
after the numbers in the second and third columns indicate the sign of
f(x) = x 3 - x - 1 at the corresponding endpoints of the intervals.
extended to solving more general equations and systems of equations. But it does
have several strong points. It always converges to a root whenever we start with an
interval whose properties are very easy to check. And it does not use derivatives
of the function f (x) as some faster methods do.
What important algorithm does the method of bisection remind you of? ffyou
have found it to closely resemble binary search, you are correct. Both of them solve
variations of the searching problem, and they are both divide-by-half algorithms.
The principal difference lies in the problem's domain: discrete for binary search
and continuous for the bisection method. Also note that while binary search
requires its input array to be sorted, the bisection method does not require its
function to be nondecreasing or nonincreasing. Finally, while binary search is very
fast, the bisection method is slow.
Method of false Position
The method offalse position (also known by its name in Latin, regula falsi) is to
interpolation search as the bisection method is to binary search. Like the bisection
method, it has, on each iteration, some interval [a,, b,] bracketing a root of a
continuous function f(x) that has opposite-sign values at a, and b,. Unlike the
bisection method, however, it computes the next root approximation not as the
middle of [a,, b,] but as the x-intercept of the straight line through the points
(a,, f(a,)) and (b,, f(b,)) (Figure 12.20).
You are asked in the exercises to show that the formula for this x-intercept
can be written as
a,J(b,)- b,.f(a,)
Xn=
f(h,)- f(a,)
(12.10)
458
Coping with the Limitations of Algorithm Power
FIGURE 12.20 Iteration of the method of false position
1
2
3
4
5
6
7
8
an
bn
Xn
f(xn)
0.00.3333330.6764710.9606191.1444251.2422591.2885321.309142-
2.0+
2.0+
2.0+
2.0+
2.0+
2.0+
2.0+
2.0+
0.333333
0.676471
0.960619
1.144425
1.242259
1.288532
1.309142
1.318071
-1.296296
-1.366909
-1.074171
-0.645561
-0.325196
-0.149163
-0.065464
-0.028173
FIGURE 12.21 Trace of the method of false position for equation (12.9). The signs
after the numbers in its second and third columns indicate the sign of
f(x) = x 3 - x -1 at the corresponding endpoints of the intervals.
EXAMPLE 2 Figure 12.21 contains the results of the first eight iterations of this
method for solving equation (12.9).
Although for this example the method of false position does not perform
as well as the bisection method, for many instances it yields a faster converging
sequence.
m
Newton's Method
Newton's method, also called the Newton-Raphson method, is one of the most
important general algorithms for solving equations. When applied to equation
(12.4) in one unknown, it can be illustrated by Figure 12.22: the next element
12.4 Algorithms for Solving Nonlinear Equations
459
FIGURE 12.22 Iteration of Newton's method
x,+l of the method's approximation sequence is obtained as the x-intercept of the
tangent line to the graph of function f (x) at x ..
The analytical formula for the elements of the approximation sequence turns
out to be
f(x,)
Xn+l = Xn - f'(xn)
for n = 0, 1, ....
(12.11)
In most cases, Newton's algorithm guarantees convergence of sequence (12.11) if
an initial approximation x 0 is chosen "close enough" to the root. (Precisely defined
prescriptions for choosing xo can be found in numerical analysis textbooks.) It may
converge for initial approximations far from the root as well, but this is not always
true.
EXAMPLE 3 Computing Fa for a 2: 0 can be done by finding a nonnegative root
2
ofequationx - a= 0. Ifweuseformnla (12.11) forthiscaseof f(x) =x 2 - a and
f'(x) = 2x, we obtain
f(x 11 )
x;- a
X 1~ +a
1
f'(x 1)
2x11
2x 11
2
Xn+l =Xn- - - =Xn- - - = - - = -(Xn
a
+ -),
X
11
which is exactly the formula we used in Section 11.4 for computing approximate
ill
values of square roots.
EXAMPLE 4 Let us apply Newton's method to equation (12.9), which we previously solved with the bisection method and the method of false position. Formula
(12.11) for this case becomes
460
Coping with the Limitations of Algorithm Power
n
x,
Xn+l
f(x,+l)
0
1
2
3
4
2.0
1.545455
1.359615
1.325801
1.324719
1.545455
1.359615
1.325801
1.324719
1.324718
1.145755
0.153705
0.004625
4.Ho- 6
5·10- 12
FIGURE 12.23 Trace of Newton's method for equation (12.9)
As an initial element of the approximation sequence, we take, say, x 0 = 2. Figure
12.23 contains the results of the first five iterations of Newton's method.
1111
You cannot fail to notice how much faster Newton's approximation sequence
converges to the root than the approximation sequences of both the bisection
method and the method of false position. This very fast convergence is typical of
Newton's method if an initial approximation is close to the equation's root. Note,
however, that on each iteration of this method we need to evaluate new values of
the function and its derivative, whereas the previous two methods require only one
new value of the function itself. Also, Newton's method does not bracket a root as
these two method' do. Moreover, for an arbitrary function and arbitrarily chosen
initial approximation, its approximation sequence may diverge. And, because
formula (12.11) has the function's derivative in the denominator, the method may
break down if it is equal to zero. In fact, Newton's method is most effective when
J'(x) is bounded away from zero near root x'. In particular, if
1/'(x)l 2: m1 > 0
on the interval between x11 and x*, we can estimate the distance between x11 and
x* by using the Mean Value Theorem of calculus as follows:
f(x,)- f(x') = f'(c)(x,- x'),
where cis some point between x, and x'. Since f(x') = 0 and 1/'(c)l": m1, we
obtain
lx,- x'l ::0 lf(x,)l.
(12.12)
ml
Formula (12.12) can be used as a criterion for stopping Newton's algorithm when
its right-hand side becomes smaller than a preselected accuracy level e. Other
possible stopping criteria are
12.4 Algorithms for Solving Nonlinear Equations
461
and
l.f(x,)l < E,
where c is a small positive number. Since the last two criteria do not necessarily
imply closeness of x, to root x*, they should be considered inferior to the one
based on (12.12).
The shortcomings of Newton's method should not overshadow its principal
strengths: fast convergence for an appropriately chosen initial approximation and
applicability to much more general types of equations and systems of equations.
-----Exercises 1
1. a. Find on the Internet or in your library a procedure for finding a real root
of the general cubic equation ax 3 + bx 2 +ex + d = 0 with real coefficients.
b. What general algorithm design technique is it based on?
2. Indicate how many roots each of the following equations has.
a. xex- 1 = 0
b. x -In x = 0
c. x sin x- 1 = 0
3. a. Prove that if p(x) is a polynomial of an odd degree, then it must have at
least one real root.
b. Prove that if x 0 is a root of ann-degree polynomial p(x), the polynomial
can be factored into
p(x) = (x- x 0 )q(x),
where q (x) is a polynomial of degree n - 1. Explain what significance this
theorem has for finding roots of a polynomial.
c. Prove that if x 0 is a root of ann-degree polynomial p(x), then
p'(xo) = q(xo).
where q(x) is the quotient of the division of p(x) by x- x0 .
4. Prove inequality (12.7).
5. Apply the bisection method to find the root of the equation
3
x +x -1 =0
with an absolute error smaller than w- 2
6. Derive formula (12.10) underlying the method of false position.
7. Apply the method of false position to find the root of the equation
3
+ x -1 = 0
with an absolute error smaller than w- 2
x
462
Coping with the Limitations of Algorithm Power
8. Derive formula (12.11) underlying Newton's method.
9. Apply Newton's method to find the root of the equation
x 3 + x -1 = 0
with an absolute error smaller than Jo- 2 •
10. Give an example that shows that the approximation sequence of Newton's
method may diverge.
11. Gobbling goat There is a grassy field in the shape of a circle with a radius of
100 feet. A goat is attached by a rope to a hook at a fixed point on the field's
border. How long should the rope be to let the goat reach only half of the
grass in the field?
SUMMARY
"
Backtracking and branch-and-bound are two algorithm design techniques for
solving problems in which the number of choices grows at least exponentially
with their instance size. Both techniques construct a solution one component
at a time, trying to tenninate the process as soon as one can ascertain that no
solution can be obtained as a result of the choices already made. This approach
makes it possible to solve many large instances of NP-hard problems in an
acceptable amount of time.
"
Both backtracking and branch-and-bound employ, as their principal mechanism, a state-space tree-a rooted tree' whose nodes represent partially
constructed solutions to the problem in question. Both techniques terminate
a node as soon as it can he guaranteed that no solution to the problem can be
obtained by considering choices that correspond to the node's descendants.
"
Backtracking constructs its state-space tree in the depth-first search fashion
in the majority of its applications. If the sequence of choices represented by a
current node of the state-space tree can be developed further without violating
the problem's constraints, it is done by considering the first remaining
legitimate option for the next component. Otherwise, the method backtracks
by undoing the last component of the partially built solution and replaces it
by the next alternative.
"
Branch-and-bound is an algorithm design technique that enhances the idea
of generating a state-space tree with the idea of estimating the best value
obtainable from a current node of the decision tree: if such an estimate is not
superior to the best solution seen up to that point in the processing, the node
is eliminated from further consideration.
Summary
463
Ill
Approximation algorithms are often used to find approximate solutions to
difficult problems of combinatorial optimization. The performance ratio is the
principal metric for measuring the accuracy of such approximation algorithms.
Ill
The nearest-neighbor and multifragment heuristic are two simple greedy
algorithms for approximating a solution to the traveling salesman problem.
The performance ratios of these algorithms are unbounded above, even for
the important subset of Euclidean graphs.
"'
The twice-around-the-tree and Christophides algorithms exploit the graph's
minimum spanning tree to construct an Eulerian circuit and then transform it
into a Hamiltonian circuit (an approximate solution to the TSP) by shortcuts.
For Euclidean graphs, the performance ratios of these algorithms are 2 and
1.5, respectively.
.,
Local search heuristics-the 2-opt, 3-opt, and Lin-Kernighan algorithmswork by replacing a few edges in the current tour to find a shorter one until
no such replacement can be found. These algorithms are capable of finding
in seconds a tour that is within a few percent of optimum for large Euclidean
instances of the traveling salesman problem.
"'
A sensible greedy algorithm for the knapsack problem is based on processing
an input's items in descending order of their value-to-weight ratios. For the
continuous version of the problem, this algorithm always yields an exact
optimal solution.
"'
Polynomial approximation schemes for the knapsack problem are polynomialtime parametric algorithms that approximate solutions with any predefined
accuracy level.
II
Solving nonlinear equations is one of the most important areas of numerical
analysis. While there are no formulas for roots of nonlinear equations (with
a few exceptions), several algorithms can solve them approximately.
.,
The bisection method and the method of false position are continuous
analogues of binary search and interpolation search, respectively. Their
principal advantage lies in bracketing a root on each iteration of the algorithm.
m
Newton's method generates a sequence of root approximations that are
x-intercepts of tangent lines to the function's graph. With a good initial
approximation, it typically requires just a few iterations to obtain a highaccuracy approximation to the equation's root.
,jl l I
1
,: I '
IIt : ,
I'
I
i:
i
i
Epilogue
Science is nothing but trained and organized common sense.
-Thomas H. Huxley {1825-18951, English biologist and educator
Well, we have arrived. It was a long road. Not as long as it took humanity to travel
from Euclid's algorithm, which opens this book, to the latest developments in
algorithmics, but it was long enough. So let us now take a last look at what we
have learned during the journey.
We started with the well-accepted proposition that the notion of an algorithm
constitutes the cornerstone of computer science. And since computer programs
are just implementations of algorithms on particular machines, algorithms lie at
the heart of practical computing, too.
Like any science, computer science is concerned with classifying its principal
subject. Although algorithms can be classified in numerous ways, two of them
are particularly important. We can classify algorithms by their underlying design
technique and by their efficiency. These two principal dimensions reflect the needs
of computing practice as well: we need design techniques as a guide for developing
a new algorithm, and we need a framework to ascertain the efficiency of a given
algorithm.
We discussed ten general design techniques in this book:
brute force
greedy technique
divide-and-conquer
decrease-and-conquer
dynamic programming
iterative improvement
transform-and-conquer
.space-time tradeoffs
branch-and-bound
backtracking
We showed how these techniques apply to a variety of important problems in
computer science: sorting, searching, string processing, graphs, and some geomet-
ric and numerical problems. Although these basic techniques are not all applicable
465
466
Epilogue
to every problem, taken collectively they provide a powerful tool kit for designing
new algorithms and classifying existing ones. Moreover, these techniques can be
thought of as general problem-solving approaches not limited to the computing
domain. The puzzles included in the book make this point.
The analysis framework classifies algorithms by the order of growth of their
running time as a function of input size. It does so by investigating the number of
times the algorithm's basic operation is executed. The main tools are summation
formulas and recurrence relations for nonrecursive and recursive algorithms, re-
spectively. We saw that a surprisingly large number of algorithms fall into one of
the few classes on the following list.
important examples
class
notation
constant time
logarithmic
8(1)
hashing (on average)
binary search (worst and average cases)
advanced sorting algorithms
cubic
El(log n)
El(n)
El(nlogn)
El(n 2)
El(n 3)
exponential
Q(2")
combinatorial problems
linear
"n-Iog-n"
quadratic
sequential search (worst and average cases)
elementary sorting algorithms
Gaussian elimination
For some algorithms, we must distinguish between the worst-, best-, and
average-case efficiencies. The average case is particularly difficult to investigate,
and we discussed how one can do this empirically.
We touched on the limitations of algorithms. We saw that there are two major
reasons for such limitations: the intrinsic complexity of a problem and the necessity
of dealing with round-off numbers for most numerical problems. We also discussed
approaches for dealing with such limitations.
It should come as no surprise, however, that there are areas of algorithmics
that were not covered in this book. The most important of them are randomized
and parallel algorithms. A randomized algorithm is an algorithm that makes random choices during its execution. For example, we can randomly select an array's
element to serve as a pivot in sorting the array by quicksort. Unlike a deterministic algorithm, a randomized algorithm behaves differently on different runs with
the same input and may even yield different results. For many applications, this
variability can be to our advantage, yielding randomized algorithms that are faster
or simpler (or both) than their deterministic counterparts.
One of the most impressive randomized algorithms discovered to date is the
Miller-Rabin algorithm for primality testing of integers (e.g., [CorOl]). This algorithm solves the problem in an acceptable amount of time for thousand-digit
numbers with the probability of yielding an erroneous answer smaller than the
probability of hardware malfunction. It is much faster than the best-known deterministic algorithms for solving this problem, which is crucial for modern cryp-
Epilogue
467
tology. If you want to learn more about raudomized algoritluns, the mouograph
by R. Motwani and P. Raghavan [Mot95] and the excellent survey by R. M. Karp
[Kar91] are particularly recommended.
The vast majority of computers in use today still closely resemble the machine outlined more than half a century ago by John von Neumann. The central
assumption of this architecture is that instructions are executed one after another,
one operation at a time. Accordingly, algorithms designed to be executed on such
machines are called sequential algorithms. It is this kind of algorithm that we
discussed in this book. The central assumption of the von Neumann model does
not hold for some newer computers that can execute operations concurrently, i.e.,
in parallel. Algorithms that take advantage of this capability are called parallel
algorithms.
Consider, as an example, the problem of computing the sum of n numbers
stored in an array A[O .. n- 1]. One can prove that any sequential algorithm that
uses only multiplications, additions, and subtractions requires at least n - 1 steps
to solve this problem. However, if we can pair and find the sum of elements A [OJ
and A[1 ], A[2] and A[3], and so on in parallel, the size of the problem will be
reduced by half. Repeating this operation until the entire sum is computed yields
an algorithm that requires just rlog2 11 l steps.
There is a wealth of books devoted to parallel algorithms. Many generalpurpose algorithm textbooks include separate chapters on parallel algorithms
([Hor98] provides particularly extensive coverage of them) or discuss them together with sequential algorithms (see [Ber04], [Mi!OO]).
The juggernaut of technological innovations has recently produced some
promising breakthroughs-such as quantum computing and DNA computingthat might have a dramatic impact on the computational capabilities and algorithms ofthe future. Quantum computing (see·[NCOO]) seeks to exploit a quantum
physics phenomenon of having an atom in two states simultaneously. Hence, at
least theoretically, a system of n such atoms, called "qubits," can simultaneously
contain 2" bits of information. In 1994, Peter Shor of AT&T Research Labs presented an algorithm for factoring an integer that took advantage of this theoretical
possibility [Sho94]. Moreover, IBM researchers were able to build a seven-qubit
computer that actually implemented Shor's algorithm and successfully factored
the number 15 into 3 and 5. Although technological problems of scaling this and
similar approaches to larger problems are formidable and may still prove to be
insurmountable, quantum computing has the potential to change our current idea
of intractable problems.
If quantum computing seeks to harness the power of quantum physics to solve
difficult computational problems, DNA computing seeks to accomplish the same
goal via exploiting the mechanism of gene selection. The most famous example
of this approach was provided in the same year, 1994, by Len Adlcman [Adl94],
a U.S. computer scientist well-known for his participation in the development of
the very important RSA encryption algorithm. (A good description and analysis of Adleman's experiment can be found in [BaaOO], Section 13.9.) Adleman
468
Epilogue
showed how the problem of finding a Hamiltonian path in a directed graph can
in principle be solved by generating DNA strands representing paths in the graph
and discarding those of them that did not satisfy the definition of such a path. The
existence of a Hamiltonian path is known to be an NP-complete problem, and
Adleman's approach is similar to exhaustive search. But a huge number of biochemical processes are occurring in parallel, leaving the hope of yielding a solution
in an acceptable amount of time. Adleman was able to solve the Hamiltonian path
problem for a small graph of seven vertices, although he had to repeat parts of
the procedure several times to purify the DNA solution. Scaling Adleman's approach to larger graphs would require an exponentially fast -growing number of
the nucleotides needed by this procedure. Although the true potential of DNA
computing remains unclear, several research teams around the world continue
working on it. Among a few reported successes, NASA announced in 2002 that a
team led by Adleman developed a DNA computer that solved a problem requiring evaluation of one million alternatives for their ability to satisfy 24 different
criteria.
So, whichever direction you take in your future journey through the land of
algorithms in your studies and your career, the road ahead is as exciting as it has
ever been. Not many areas of science and engineering can make this claim with
the same assurance that algorithmics can. Have a great trip!
APPENDIX A
Useful Formulas for the
Analysis of Algorithms
This appendix contains a list of useful formulas and rules that are helpful in the
mathematical analysis of algorithms. More advanced material can be found in
[Gre82], [Gra94], [Pur85], and [Sed96].
Properties of Logarithms
All logarithm bases are assumed to be greater than 1 in the formulas below; lg x
denotes the logarithm base 2, In x denotes the logarithm base e = 2.71828 ... ;
x, yare arbitrary positive numbers.
1.
log" 1 = 0
2.
log" a= 1
3.
log, xY = y log, x
4.
log, xy = log" x + log, y
5.
log, - = log" x - log, y
7.
Jog, x = - - = log, b Iogb x
X
y
Iogbx
Iogb a
Combinatorics
1.
Number of permutations of ann-element set: P(n) = n!
2.
Number of k-combinations of ann-element set: C(n, k) = k!(nn! k)!
3.
Number of subsets of ann-element set: 2"
469
470
Useful Formulas for the Analysis of Algorithms
Important Summation Formulas
1.
1
3.
4.
L 1=n
i=l
i=l
u-1+1 times
~.
n(n + 1)
id
2
~r=1+2+ ... +n=
t
i~l
1 2
~-n
2
;2 = 12 + 22 + ... + 11 2 = n(n + 1)(2n + 1) ~ ~
6
11 3
3
1
Ln ;k = 1k + 2k + ... + "k ~--nk+1
k+l
id
n
5.
n
L" 1 = ~ =u -I+ 1 (1. u are integer limits,/<: u);
L ai
=
an+l ~ 1
----c- (a ofo 1);
1 +a + · · · + a"
a-1
i=O
n
Li=2n+l_l
i=O
n
6.
L i2' = 1 * 2 + 2 * 22 + ... + n2" = en - 1)2"+1 + 2
i=l
7.
L" 1~ = 1 + -1 + · · · + -1 ~ ln n + y, where y ~ 0.5772 ... (Euler's constant)
i=l
2
r
n
n
8.
Llgi~nlgn
i=l
Sum Manipulation Rules
1.
II
II
i=l
i=l
Lea; =c La
u
2.
u
4.
u
Lea, ±b;)= La,± Lb,
i=l
i=l
i=l
3.
1
II
Ill
i=l
i=l
II
La;= L a+ L a;, where l ::=: m u
<
1
i=m+-1
"
Lcaia;-t) =au- al~l
i=l
Useful Formulas for the Analysis of Algorithms
471
Approximation of a Sum by a Definite Integral
u
u
J
~ f(i) .:"
I
i=l
_ f(x)dx .:"
1 1
fu+l
f(x)dx for a nondecreasing f(x)
1
f u+l f(x)dx .:"I:u f(i) .:" j" f(x)dx for a nonincreasing f(x)
/-1
Floor and Ceiling Formulas
The floor of a real number x, denoted [xj, is defined as the greatest integer not
larger than x (e.g., [3.8j = 3, [-3.8j = -4, [3j = 3). The ceiling of a real number
X, denoted fX 1, is defined as the Smallest integer not smaller than X (e.g., f3.81 = 4,
r-3.81 = -3, f31 = 3).
1,
X- 1 < [xj .:"X.:" fx1 <X+ 1
2.
3.
[x+nJ = [xj +nand fx+n1 = fx1 +n forrealxandintegern
[n/2J + rn/21 = n
4.
flg(n + 1)1 = [lg nj + 1
Miscellaneous
~ ( ~)" as n -+ oo (Stirling's formula)
1.
n! "'
2.
Modular arithmetic (n, mare integers, pis a positive integer)
(n + m) modp = (n modp +m mod p) mod p
(n
* m) mod p = ((n mod p) * (m mod p)) mod p
~~.' ·i:'
.• I
Ii,
Ii
APPENDIX
Short Tutorial on
Recurrence Relations
Sequences and Recurrence Relations
DEFINITION
A (numerical) sequence is an ordered list of numbers.
Examples: 2, 4, 6, 8, 10, 12, ... (positive even integers)
0, 1, 1, 2, 3, 5, 8, ... (the Fibonacci numbers)
0, 1, 3, 6, 10, 15, ... (numbers of key comparisons in selection sort)
A sequence is usually denoted by a letter (such asx or a) with a subindex (such
as nor i) written in curly brackets, e.g., {x,}. We use the alternative notationx(n).
Tbis notation stresses the fact that a sequence is a function: its argument n indicates
a position of a number in the list, while the function's value x(n) stands for that
number itself. x (n) is called the generic term of the sequence.
There are two principal ways to define a sequence:
"'
by an explicit formula expressing its generic term as a function of n, e.g.,
x(n) = 2n for n 2: 0
!!ill
by an equation relating its generic term to one or more other terms of the
sequence, combined with one or more explicit values for the first term(s),
e.g.,
x(n) = x(n- 1)
x(O) = 0
+n
for n > 0
(8.1)
(B.2)
It is the latter method that is particularly important for analysis of recursive
algorithms (see Section 2.4 for a detailed discussion of this topic).
An equation such as (B1) is called a recurrence equation or recurrence relation
(or simply a recurrence), while an equation such as (B2) is called its initial
473
474
Short Tutorial on Recurrence Relations
condition. An initial condition can be given for a value of n other than 0 (e.g., for
n = 1) and for some recurrences (e.g., for recurrence F(n) = F(n -1) + F(n- 2)
defining the Fibonacci numbers-see Section 2.5), more than one value needs to
be specified by initial conditions.
To solve a given recurrence subject to a given initial condition means to find
an explicit formula for the generic term of the sequence that satisfies both the
recurrence equation and the initial condition or to prove that such a sequence does
not exist. For example, the solution to recurrence (B.1) subject to initial condition
(B.2) is
n ("-'n--'+'--1='-)
2
x (n) = :.:..
for n 2: 0.
(B.3)
It can be verified by substituting this formula into (B.1) to check that the equality
holds for every n > 0, i.e., that
n(n + 1)
2
=
(n -1)(n -1+ 1)
2
+n ,
and into (B.2) to check that x(O) = 0, i.e., that
0(0+1) =0.
2
Sometimes it is convenient to distinguish between a general solution and
a particular solution to a recurrence. Recurrence equations typically have an
infinite number of sequences that satisfy them. A general solution to a recurrence
equation is a formula that specifies all such sequences. Typically, a general solution
involves one or more arbitrary constants. For example, for recurrence (B.J ), the
general solution can be specified by the for111ula
x(n)=c+ n(n:1),
(B.4)
where cis such an arbitrary constant. By assigning different values to c, we can
get all the solutions to equation (B.1) and only these solutions.
A particular solution is a specific sequence that satisfies a given recurrence
equation. Usually we are interested in a particular solution that satisfies a given
initial condition. For example, sequence (B.3) is a particular solution to (B.J )(B.2).
Methods for Solving Recurrence Relations
No universal method exists that would enable us to solve every recurrence relation. (This is not surprising, because we do not have such a method even for solving
much simpler equations in one unknown f(x) = 0 for an arbitrary function f(x).)
There are several techniques, however, some more powerful than others, that can
solve a variety of recurrences.
Short Tuta rial on Recurrence Relations
475
Method of forward substitutions Starting with the initial term (or terms) of the
sequence given hy the initial condition(s), we can use the recurrence equation to
generate the first few terms of its solution in the hope of seeing a pattern that can be
expressed by a closed-end formula. If such a formula isJound, its validity should be
either checked by direct substitution into the recurrence equation and the initial
condition (as we did for (B.1)-(B.2)) or proved by mathematical induction.
For example, consider the recurrence
x(n) = 2x(n -
1) + 1
for n > 1
x(l) = 1.
(B.S)
(B.6)
We obtain the first few terms as follows:
x(1) = 1
x (2) = 2x (1) + 1 = 2 · 1 + 1 = 3
x(3) = 2x(2) + 1 = 2 · 3 + 1 = 7
x(4) = 2x(3) + 1 = 2 · 7 + 1 = 15.
It is not difficult to note that these numbers are one less than consecutive powers
of2:
x(n) = 2"- 1
for 11 = 1, 2, 3, and 4.
We can prove the hypothesis that this formula yields the generic term of the
solution to (B.5)-(B.6) either by direct substitution of the formula into (B.5) and
(B.6) or by mathematical induction.
As a practical matter, the method of forward substitutions works in a very
limited number of cases because it is usually very difficult to recognize the pattern
in the first few terms of the sequence.
Method of backward substitutions This method of solving recurrence relations
works exactly as its name implies: using the recurrence relation in question, we
express x(n - 1) as a function of x(n- 2) and substitute the result into the original
equation to get x(n) as a function of x(n- 2). Repeating this step for x(112) yields an expression of x (n) as a function of x (11 - 3). For many recurrence
relations, we will then be able to see a pattern and express x(n) as a function of
x(n - i) for an arbitrary i = 1, 2, .... Selecting i to make 11 - i reach the initial
condition and using one of the standard summation formulas often leads to a
closed-end formula for the solution to the recurrence.
As an example, let us apply the method of backward substitutions to recurrence (B.1)-(B.2). Thus, we have the recurrence equation
x(11) = x(n- 1) + 11.
Replacing n by n- 1 in the equation yields x(n- 1) = x(11- 2) + n- 1; after
substituting this expression for x(11 - 1) in the initial equation, we obtain
x(11) = [x(n - 2)
+ 11 - 1J + 11 = x(n - 2) + (n - 1) + n.
ShortTutorial on Recurrence Relations
476
Replacingn byn- 2intheinitialequationyieldsx(n- 2) =x(n- 3) +n- 2;after
substituting this expression for x(n - 2), we obtain
x(n) = [x(n- 3) + n- 2] + (n- 1) + n = x(n- 3) + (n- 2) + (n- 1) + n.
Comparing the three formulas for x(n), we can see the pattern that arises after i
such substitutions: 1
x(n) = x(n- i) + (n- i + 1) + (n- i + 2) + · · · + n.
Since initial condition (B.2) is specified for n = 0, we need n - i = 0, i.e., i = n, to
reach it:
x(n) = x(O) + 1 + 2 + · · · + n = 0 + 1 + 2 + · · · + n = n(n + 1)/2.
The method of backward substitutions works surprisingly well for a wide
variety of simple recurrence relations. You can find many examples of its successful
applications throughout this book (see, in particular, Section 2.4 and its exercises).
Linear second-order recnrrences with constant coefficients An important class
of recurrences that can be solved by neither forward nor backward substitutions
are recurrences of the type
ax(n) + bx(n- 1) + cx(n- 2) = .f(n),
(B.7)
where a, b, and care real numbers, a oJ 0. Such a recurrence is called second-order
linear recurrence with constant coefficients. It is second-order because elements
x(n) and x(n- 2) are two positions apart in the unknown sequence in question; it
is linear because the left -hand side is a linear combination of the unknown terms
of the sequence; it has constant coefficients because of the assumption that a, b,
and care some fixed numbers. If .f(n) = 0 for every n, the recurrence is said to be
homogeneous; otherwise, it is called inhomogeneous.
Let us consider first the homogeneous case:
ax(n) + bx(n -1) + cx(n- 2) = 0.
(B.S)
Except for the degenerate situation of b = c = 0, equation (B.8) has intinitelymany
solutions. All these solutions, which make up the general solution to (B.8), can be
obtained by one of the three formulas that follow. Which of the three formulas
applies to a particular case depends on the roots of the quadratic equation with
the same coefficients as recurrence (B.8):
ar 2 + br + c = 0.
1.
(B.9)
Strictly speaking, the validity of the pattern's formula needs to be proved by mathematical induction
on i. It is often easier, however, to get the solution first and then verify it (e.g., as we did earlier for
x(n) = n(n + 1)/2).
Short Tutorial on Recurrence Relations
477
Quadratic equation (B.9) is called the characteristic equation for recurrence
equation (B.S).
THEOREM 1 Let r1• r 2 be two roots of characteristic equation (B.9) for recurrence relation (B.8).
Case 1 If r 1 and r 2 are real and distinct, the geueral solution to recurrence
(B.8) is obtained by the formula
x(n) = ar~ + f3r~1 ,
1
where a and fJ are two arbitrary real constants.
Case 2 If r1 and rz are equal to each other, the general solution to recurrence
(B.8) is obtained by the formula
x(n) = ar 11 + f3nr 11 ,
where r = r 1 = r 2 and a and fJ are two arbitrary real constants.
Case 3 If r1 2 = u ± i v are two distinct complex numbers, the general solution
to recurrence (B.8) is obtained as
x(n) = y"[a cos n8 + fJ sin nB],
2
2
where y = J u + v • 8 =arctan vju, and a and fJ are two arbitrary real constants.
Case 1 of this theorem arises, in particular, in deriving the explicit formula
for the nth Fibonacci number (Section 2.5). As another example, let us solve the
recurrence
x(n)- 6x(n- 1) + 9x(n- 2) = 0.
Its characteristic equation
has two equal roots r1 = rz = 3. Hence, according to Case 2 of Theorem 1, its
general solution is given by the formula
x(n) = a3" + fJn3".
If we want to fmd its particular solution for which, say, x(O) = 0 and x(l) =
3, we substitute n = 0 and n = 1 into the last equation to get a system of two
linear equations in two unknowns. Its solution is a = 0 and fJ = 1, and hence the
particular solution is
x(n)=n3".
Let us now turn to the case of inhomogeneous linear second-order recurrences
with constant coefficients.
478
Short Tutorial on Recurrence Relations
THEOREM 2 The general solution to inhomogeneous equation (B.7) can be
obtained as the sum of the general solution to the corresponding homogeneous
equation (B.S) and a particular solution to inhomogeneous equation (B.7).
Since Theorem 1 gives a complete recipe for finding the general solution to a
homogeneous second-order linear equation with constant coefficients, Theorem
2 reduces the task of finding all solutions to equation (B.7) to finding just one
particular solution to it. For an arbitrary function f(n) in the right-hand side
of equation (B.7), it is still a difficult task with no general help available. For
a few simple classes of functions, however, a particular solution can be found.
Specifically, if f (n) is a nonzero constant, we can look for a particular solution
that is a constant as well.
As an example, let us find the general solution to the inhomogeneous recurrence
x(n)- 6x(n- 1) + 9x(n- 2) = 4.
If x (n) = c is its particular solution, constant c must satisfy the equation
c- 6c+9c=4,
which yields c = 1. Since we have already found the general solution to the corresponding homogeneous equation
x(n)- 6x(n- 1) + 9x(n- 2) = 0,
the general solution to x (n) - 6x (n - 1) + 9x (n - 2) = 4 is obtained by the formula
x(n) = a3n + {Jn3n + 1.
Before leaving this topic, we should note that the results analogous to those of
Theorems 1 and 2 hold for the genera/linear k-th degree recurrence with constant
coefficients
akx(n) + ak-lx(n- 1) + · · · + aox(n- k) = f(n).
(B.lO)
The practicality of this generalization is limited, however, by the necessity of
finding roots of the k-th degree polynomial
(B.l1)
which is the characteristic equation for recurrence (B.10).
Finally, there are several other, more sophisticated techniques for solving
recurrence relations. Purdom and Brown [Pur85] provide a particularly thorough
discussion of this topic from the analysis of algorithms perspective.
Short Tutorial on Recurrence Relations
479
Common Recurrence Types in Algorithm Analysis
There are a few recurrence types that arise in the analysis of algorithms with
remarkable regularity. This happens because they reflect one of the fundamental
design techniques.
Decrease-by-one A decrease-by-one algorithm solves a problem by exploiting a
relationship between a given instance of size n and a smaller instance of size n - 1.
Specific examples include recursive evaluation of n! (Section 2.4) and insertion sort
(Section 5.1). The recurrence equation for investigating the time efficiency of such
algorithms typically has the form
T(n) = T(n -1) + f(n),
(B.l2)
where function f (n) accounts for the time needed to reduce an instance to a
smaller one and to extend the solution of the smaller instance to a solution of
the larger instance. Applying backward substitutions to (B.12) yields
T(n) = T(n- 1) + f(n)
= T (n - 2)
= T(O) +
+ f (n - 1) + f (n)
L" f(j).
j=l
For a specific function .f(x), the sum I:'jd f(j) can usually be either computed
exactly or its order of growth ascertained. For example, if .f(n) = 1, I:]~ 1 .f(j) = n;
if .f(n) =log n, I:}~dUJ E E>(n log n); if .f(n) = nk, I:}~d(j) E E>(nk+ 1). The
sum I:}~r .f(j) can also be approximated by formulas involving integrals (see, in
particular, the appropriate formulas in Appendix A).
Decrease-by-a-constant-factor A decrease-by-a-constant-factor algorithm
solves a problem by reducing its instance of size n to an instance of size n/b (b = 2
for most but not all such algorithms), solving the smaller instance recursively, and
then, if necessary, extending the solution of the smaller instance to a solution of
the given instance. The most important example is binary search; other examples
include exponentiation by squaring (introduction to Chapter 5), multiplication ii
la russe, and the fake-coin problem (Section 5.5).
The recurrence equation for investigating the time efficiency of such algorithms typically has the form
T(n) = T(n/b) + .f(n),
(B.13)
where b > 1 and function f (n) accounts for the time needed to reduce an instance
to a smaller one and to extend the solution of the smaller instance to a solution
of the larger instance. Strictly speaking, equation (B.l3) is valid only for n = bk,
Short Tutorial on Recurrence Relations
480
k = 0, 1, .... For values ofn that are not powers of b, there is typically some roundoff, usually involving the floor and/or ceiling functions. The standard approach to
such equations is to solve them for n = bk first. Afterward, either the solution is
tweaked to make it valid for all n 's (see, for example, Section 4.3 and Problem 3 in
Exercises 4.3), or the order of growth of the solution is established based on the
smoothness rule (Theorem 4 in this appendix).
By considering n = bk, k = 0, 1, ... , and applying backward substitutions to
(B.13), we obtain the following:
T(b') = T(bk-l) + f(b')
= T(bk-2) + f(bk-1) + f(b')
k
= T(1) +
L f(bj).
)=1
For a specific function f(x), the sum L~d f(bj) can usually be either computed
exactly or its order of growth ascertained. For example, if f (n) = 1,
k
L f(bj) = k = Iogb n.
j=l
If f (n) = n, to give another example,
k
k
k
L:tc"jl='L"j=""
-1="n-1
j~l
j~1
b- 1
b- 1
Also, recurrence (B.13) is a special case of recurrence (B.14) covered by the Master
Theorem (Theorem 5 in this appendix). According to this theorem, in particular,
if f(n) E Q(nd) where d > 0, then T(n) E Q(nd) as well.
Divide-and-conquer A divide-and-conquer algorithm solves a problem by dividing its given instance into several smaller instances, solving each of them recursively, and then, if necessary, combining the solutions to the smaller instances into
a solution to the given instance. Assuming that all smaller instances have the same
size n/ b, with a of them being actually solved, we get thefollowing recurrence valid
for n = b', k = 1, 2, ... :
T(n) = aT(n/b) + f(n),
(B.14)
where a 2: 1, b 2: 2, and f(n) is a function that accounts for the time spent on
dividing the problem into smaller ones and combining their solutions. Recurrence
(B.J 4) is called the general divide-and-conquer recurrence?
2.
In our terminology, for a= 1, it covers decrease-by-a-constant-factor, not divide-and-conquer, algorithms.
T
.
.
l
Short Tutorial on Recurrence Relations
481
Applying backward substitutions to (B.14) yields the following:
T(b') = aT(b'- 1) + f(b')
= a[aT(bk-2) + f(bk-l)] + f(b') = a2T(bk-2) + af(bk-1) + f(b')
= a2[aT(bk-3) + f(l}-2)] + af(bk-l) + f(b')
= a3T(bk-3)
= a'T(l)
+ a2 f(bk-2) + af(bk-l) + f(b')
+ ak-1 f(b1) + ak-2 f(b2) + ... +a" f(b')
k
= a'[T(l) +
L f(bi)/ai].
j=l
Since ak =a 1°gb 11 = nlog, a, we get the following formula for the solution to recur-- renee (B.14) for n = h"
log"1l
T(n) = n!og,a[T(l) +
L f(bi)/ai].
(JUS)
J=l
Obviously, the order of growth of solution T (n) depends on the values of the
constants a and b and the order of growth of the function f (n). Under certain
assumptions about f(n) discussed in the next section, we can simplify formula
(B.15) and get explicit results about the order of growth of T(n).
Smoothness Rule and the Master Theorem We mentioned earlier that the time
efficiency of decrease-by-a-constant-factor and divide-and-conquer algorithms is
usually investigated first for n's that are powers of b. (Most often, b = 2, as it is
in binary search and mergesort; sometimes b = 3, as it is in the better algorithm
for the fake-coin problem of Section 5.5, but it can be any integer greater than or
equal to 2.) The question we are going to address now is when the order of growth
observed for n's that are powers of b can be extended to all its values.
DEFINITION Let f(n) be a nonnegative function defined on the set of natural
numbers. f (n) is called eventually nondecreasing if there exists some nonnegative
integer n 0 so that f(n) is nondecreasing on the interval [n 0 , oo), i.e.,
For example, the function (n - 100/ is eventually non decreasing, although it
is decreasing on the interval [0, 100], and the function sin2 rr;• is a function that is
not eventually nondecreasing. The vast majority of functions we encounter in the
analysis of algorithms are eventually nondecreasing. (Most of them are, in fact,
nondecreasing on their entire domains.)
I
li
H
I'
ii
''
I:
482
Short Tutorial on Recurrence Relations
DEFINITION Let 1(11) be a nonnegative function defined on the set of natural
numbers. I (n) is called smooth if it is eventually nondecreasing and
1(2n) E 8(f(n)).
It is easy to check that functions which do not grow too fast, including log 11,
11, n log 11, and 11" where a 2:0, are smooth. For example, l(n) = n log 11 is smooth
because
1(2n) = 2n log 2n = 2n(log 2 +log n) = (2log 2)n + 211log n E C0(n log n).
Fast-growing functions, such as a 11 where a> 1 and n!, are not smooth. For example, f(n) = 2" is not smooth because
f(2n) = 22" = 4" 'f 8(2").
THEOREM 3
Let I (n) be a smooth function as just defined. Then, for any fixed
integer b 2: 2,
f(bn) E 8(/(n)),
i.e., there exist positive constants cb and db and a nonnegative integer n 0 such that
d 1,f(n) :0: .f(bn) :0: cbl(n)
for n 2:110.
(The same assertion, with obvious changes, holds for the 0 and S1 notations.)
PROOF We will prove the theorem for the 0 notation only; the proof of the S1
part is the same. First, it is easy to check by induction that if f (2n) :0: c2f (n) for
n :::::_ n0 , then
1(2kn) :0: c~l(n)
fork= 1, 2, ... and n 2:110.
The induction basis for k = 1 checks out trivially. For the general case, assuming
that 1(2k- 1n) :0: c~-l l(n) for n 2: 11 0, we obtain
.f(2kn) = .f(2 · zk-Jn) :": Czl(2k··ln) :": CzC~-l.f(n) = c~f(n).
(This proves the theorem for h = zk.) Consider now an arbitrary integer b 2:2. Let
k be a positive integer such that zk-t :s h < zk We can estimate .f(bn) above by
assuming without loss of generality that I (n) is nondccreasing for n 2: n 0:
f(bn) :0: .f(2kn) :0: c~.f(n).
Hence, we can use c~ as a required constant for this value of h to complete the
proof.
Ill
Short Tutorial on Recurrence Relations
483
The importance of the notions introduced above stems from the following
theorem.
THEOREM 4 (Smoothness Rule) Let T(n) be an eventually nondecreasing
function and f (n) be a smooth function. If
T(n) E 8(.f(n)) for values of n that are powers of b,
where b 2: 2, then
T(n) E 8(.f(n)).
(The analogous results hold for the cases of 0 and Q as well.)
PROOF We will prove just the 0 part; the Q part can be proved by the analogous
argument. By the theorem's assumption, there exist a positive constant c and a
positive integer n 0 = bko such that
T(n) is nondecreasing for n 2: n 0 , and f(bn) S cbf(n) for n 2: n 0 by Theorem 3.
Consider an arbitrary value of n, n 2: n0 . It is bracketed by two consecutive powers
of b: n 0 S bk :" n < bk+ 1 Therefore,
Hence, we can use the product ccb as a constant required by the 0 (f (n)) dcfmition
to complete the 0 part of the theorem's proof.
!!I
Theorem 4 allows us to expand the information about the order of growth
established for T(n) on a convenient subset of values (powers of b) to its entire
domain. Here is one of the most useful assertions of this kind.
THEOREM 5 (Master Theorem) Let T (n) be an eventually nondecreasing
function that satisfies the recurrence
T(n)
= aT(njb) + f(n)
for n = bk, k = l, 2, ...
T(l) = c,
where a 2: l, b 2: 2, c > 0. If f(n) E 8(n"J where d 2: 0, then
if
a < b"
8(n" log n)
if
E>(n 10gb a)
if
a= hd
a > bd.
E>(n")
T(n) E
{
(Similar results hold for the 0 and Q notations, too.)
484
Short Tutorial on Recurrence Relations
PROOF We will prove the theorem for the principal special case off (n) = nd. (A
proof of the general case is a minor technical extension of the same argumentsee, e.g., [CorOl].) If f(n) = nd, equality (B.15) yields for n = bk, k = 0, 1, ... ,
T(n) = n 10g""[T(l) +
Iogb n
\ogh n
}=1
)=1
L bjd laj] = nlog,a[T(1) + L (bd la)j].
The sum in this formula is that of a geometric series, and therefore
and
Iogbn
L (bd la)j = Iogb n if bd =a.
}=1
If a < bd, then bd 1a > 1, and therefore
Hence,
1ogb 11
T(n) =n 1"g""[T(1) +
L (bdla)j]
E
n 10 gba8((b"la) 10 g"n)
j=l
= 8(nlog,a(bd la)log,n) = S(alogbn(bd la)logbn)
= 8(bd log,") = i,<)(blog, "") = 8(n").
If a > bd, then b" I a < 1, and therefore
Iogb n
T(n) = nlog, "[T(1) +
L (hd la)j] E 8(n gh ").
10
}=1
Short Tutorial on Recurrence Relations
485
If a = bd, then bd j a = 1, and therefore
logbn
T(n) =nlog,a[T(l) +
L (hdja)j]=n g""[T(1) +logbn]
10
}=l
E 8(n
10
g" a 1ogb n) = 8(n 10g" bd 1ogb n) = El(nd 1ogb n).
Since f (n) = nd is a smooth function for any d, a reference to Theorem 4 completes
1111
the proof.
Theorem 5 provides a very convenient tool for a quick efficiency analysis of
divide-and-conquer and decrease-by-a-constant-factor algorithms. You can find
examples of such applications throughout the book.
1'·...' ',11
;:
:I
&1
Bibliography
[Ade62]
Adelson-Velsky, G.M. and Landis, E.M. An algorithm for organization of information. Soviet Mathematics Doklady, val. 3, 1962,
1259-1263.
[Adl94]
Adleman, L.M. Molecular computation of solutions to combinatorial
problems. Science, vol. 266,1994,1021-1024.
[Agr02]
Agrawal, M., Kayal, N., and Saxena, N. PRIMES l' in P Preprint,
Department of Computer Science and Engineering, Indian Institute
of Technology Kanpur, Kanpur-208016, India, August 6, 2002.
[Aho74]
Abo, A.V, Hopcroft, J.E., and Ullman, J.D. The Design and Analysis
of Computer Algorithms. Addison-Wesley, Reading, MA, 1974.
[Aho83]
Abo, A. V, Hopcroft, J.E., and Ullman, J.D. Data Structures and
Algorithms. Addison-Wesley, Reading, MA, 1983.
[Ahu93]
Ahuja, R.K., Magnanti, T.L., and Oriin, J.B. Network Flows: Theory,
Algorithms, and Applications. Prentice Hall, Englewood Cliffs, NJ,
1993.
[Alb99]
Albert, R., Hawoong Jeong, and Barbiisi, A.-L. Diameter of the
World-Wide Web. Nature, vol. 401,1999,130-131.
[App]
Applegate, D., Bixby, R., Chviital, V, and Cook, W. Concord: a Code
for Solving Traveling Salesman Problems. http://www.tsp.gatcch.edu
q/concorde.html.
[Ata98]
Atallah, M.J., editor. Algorithms and Theory of Computation
Handbook. CRC Press, Boca Raton, FL, 1998.
[BaaOO]
Baase, S., and Van Gelder, A. Computer Algorithms: Introduction to
Design and Analysis, 3rd ed. Addison-Wesley, Reading, MA, 2000.
[Bae81]
Baecker, R. (with assistance of D. Sherman). Sorting Out Sorting. 30minute color sound fllm. Dynamic Graphics Project, University of
Toronto, 1981. (Distributed by Morgan Kaufmann, Publishers.)
487
i
488
Bibliography
(Bae98]
Baecker, R Sorting out sorting: a case study of software visualization
for teaching computer science. In Sofhvare Visualization: Programming as a Multimedia Experience, edited by J. Stasko, J. Domingue,
M.C Brown, and B.A. Price. MIT Press, Cambridge, MA, 1998,
369-381.
(BaY95]
Baeza-Yates, RA. Teaching Algorithms. ACM SIGACT News, vol.
26,no.4,Dec. 1995,51-59.
(Bay72]
Bayer, R, and McGreight, E.M. Organization and maintenance of
large ordered indices. Acta Informatica, vol. 1, no. 3, 1972, 173-189.
[Bel57]
Bellman, R.E. Dynamic Programming. Princeton University Press,
Princeton, NJ, 1957.
[Bel62]
Bellman, R.E., and Dreyfus, S.E. Applied Dynamic Programming.
Princeton University Press, Princeton, NJ, 1962.
(BenOO]
Bentley, J. Programming Pearls, 2nd ed. Addison-Wesley, Reading,
MA,2000.
[Ben90]
Bently, J.L. Experiments on traveling salesman heuristics. In Proceedings of the First Annual ACM-SIAM Symposium on Discrete
Algorithms, 1990, 90-99.
[Be rOO]
Berlinski, D. The Advent of the Algorithm. Harcourt, Inc., New York,
2000.
[Ber01]
Berstekas, D.P. Dynamic Programming and Optimal Control, 2nd ed.
(2 vols.). Athena Scientific, 2001.
(Ber03]
Berlekamp, E.R, Conway, J.H., and Guy, R.K. Winning Ways for
Your Mathematical Plays, 2nd ed.,,volumes 1-4. A K Peters, Natick,
MA,2003.
(Ber04]
Berman, K.A., and Paul, J.L. Algorithms: Sequential, Parallel, and
Distiributed. Course Technology, MA, 2004.
(Blo73]
Bloom, M., Floyd, RW., Pratt, V, Rivest, R.L., and Tarjan, R.E.
Time bounds for selection. Journal of Computer and System Scienc";
vol. 7, no. 4, 1973,448-461.
[Bog]
Bogomolny, A. Interactive Mathematics Miscellany and Puzzles.
http://www.cut-the-knot.org.
[Boy77]
Boyer, R.S., and Moore, J.S. A fast string searching algorithm.
Communications of the ACM, vol. 21, no. 10,1977,762-772.
(Bra96]
Brassard, G., and Bralley, P. Fundamentals of Algorithmics. Prentice
Hall, Englewood Cliffs, NJ, 1996.
[Bro84]
Brown, M., and Sedgewick, R. A system for algorithm animation.
In Proceedings of ACM SIGGRAPH '84, Minneapolis, July 1984,
177-186.
Bibliography
489
[Bro91]
Brown, M.R. ZEUS: A system for algorithm animation and multiview editing. In Proceedings of the 1991 IEEE Workshop on Visual
Languages, Kobe, Japan, October 1991,4-9.
[Bru04]
Brualdi, R. Introductory Combinatorics, 4th ed. Prentice Hall, Upper
Saddle River, NJ, 2004.
[Car79]
Carmony, L. Odd pie fights. Mathematics Teacher, vol. 72, no.1, 1979,
61-64.
[ChaOO]
Chandler, J.P. Patent protection of computer programs. Minnesota
Intellectual Property Review, vol. 1, no. 1, 2000, 33-46.
[Cha98]
Chabert, Jean-Luc, editor. A History of Algorithms: From the Pebble
to the Microchip. Translated by Chris Weeks. Springer-Verlag, Berlin,
1998.
[Chr76]
Christofides, N. Worst-case Analysis of a New Heuristic for the
Traveling Salesman Problem. Technical Report, GSIA, CarnegieMellon University, 1976.
[Chv83]
Chvatal, V. Linear Programming. W.H. Freeman, New York, NY,
1983.
[Com79]
Comer, D. The ubiquitous B-tree. ACM Computing Surveys, vol. 11,
no.2, 1979,121-137.
[Coo71]
Cook, S.A. The complexity of theorem-proving procedures. In
Proceedings of the Third Annual ACM Symposium on the Theory
of Computing, 1971,151-158.
[Coo87]
Coopersmith, D., and Winograd, S. Matrix multiplication via
arithmetic progressions. In Proceedings of Nineteenth Annual ACM
Symposium on the Theory of Computing, 1987,1-6.
Carmen, T.H., Leiserson, C.E., Rivest, R.L., and C. Stein. Introduction to Algorithms, 2nd ed. MIT Press, Cambridge, MA, 2001.
Dantzig, G.B. Linear Programming and Extensions. Princeton
University Press, Princeton, NJ, 1963.
[CarOl]
[Dan63]
[Dew93]
Dewdney, A.K. The (New) Turing Omnibus. Computer Science Press,
New York, 1993.
[Dij59]
Dijkstra, E.W. A note on two problems in connection with graphs.
Numerische Mathematik, val. 1, 1959, 269-271.
[Dud70]
Dudley, U. The first recreational mathematics book . .Journal of
Recreational Mathematics, 1970, 164-169.
Edmonds, J. Paths, trees, and flowers. Canadian .Journal of Mathematics, vol. 17,1965,449-467.
[Edm65]
[Edm72]
Edmonds, J. and Karp, R.M. Theoretical improvements in algorithmic efficiency for network flow problems. Journal of the ACM, val.
19, no. 2, 1972, 248-264.
Floyd, R.W. Algorithm 97: shortest path. Communications of the
ACM, val. 5, no. 6, 1962, 345.
Ford, L.R., Jr., and Fulkerson, D.R. A simple algorithm for finding
maximal network flows and an application to the Hitchcock problem.
Canadian .Journal of Mathematics, val. 9, no. 2, 1957, 21~218.
Ford, L.R., Jr., and Fulkerson, D.R. Flows in Networks. Princeton
University Press, Princeton, NJ, 1962.
Forsythe, G.E. What to do till the computer scientist comes. American
Mathematical Monthly, val. 75, no. 5, 1968, 454-462.
Forsythe, G.E. Solving a quadratic equation on a computer. In The
Mathematical Sciences, edited by COSRIMS and George Boehm,
MIT Press, Cambridge, MA, 1969, 138-152.
Gale, D. and Shapley, L.S. College admissions and the stability of
marriage. American Mathematical Monthly, val. 69, 1962, 9-14.
Galil, Z. Efficient algorithms for finding maximum matching in
graphs. Computing Surveys, val. 18, no. 1, March 1986, 23-38.
Gardner, M. aha! Insight. Scientific American/W.H. Freeman, New
York, NY, 1978.
Garey, M.R., and Johnson, D.S. Computers and Intractability: A
Guide to the Theory of NP-Completeness. W.H. Freeman and
Company, New York, 1979.
Gardner, M. My Best Mathematical and Logic Puzzles. Dover
Publications, New York, 1994.
Gardiner, A. Mathematical Puzzling. Dover, Mineola NY, 1999.
Gerald, C.F., and Wheatley, P.O. Applied Numerical Analysis, 7th ed.
Addison-Wesley, Reading, MA, 2003.
[Gon91]
[Goo02]
[Gra94]
Ginat, D. Embedding instructive assertions in program design. In
Proceedings of ITiCSE'04, June 23-30, 2004 Leeds, UK, 62-66.
Gloor, P.A. User interface issues for algorithm animation. In Software
Visualization: Programming as a Multimedia Experience, edited by
John Stasko, John Domingue, Marc H. Brown, and Blaine A. Price.
MIT Press, Cambridge, MA, 1998, 145-152.
Gannet, G.H., and Baeza-Yates, R. Handbook of Algorithms and
Data Structures in Pascal and C, 2nd ed. Addison-Wesley, Reading,
MA, 1991.
Goodrich, M.T., and Tamassia, R. Algorithm Design: Foundations,
Analysis, and Internet Examples. John Wiley & Sons, New York, 2002.
Graham, R.L., Knuth, D. E., and Patashnik, 0. Concrete Mathematics:
A Foundation for Computer Science, 2nd ed. Addison-Wesley,
Reading, MA, 1994.
~
1
Bibliography
I
I
1
491
[Gre82]
Green, D.H., and Knuth, D.E. Mathematics for Analysis of Algorithms, 2nd ed. Birkhauser, Boston, 1982.
[Gri81]
Gries, D. The Science of Programming. Springer-Verlag, New York,
1981.
[Gus89]
Gusfield, D. and Irwing, R.W. The Stable Marriage Problem: Structure
and Algorithms. MIT Press, 1989.
[Gut02]
Gutin, G. and Punnen, A.P., eds. Traveling Salesman Problem and Its
Variations. Kluwer Academic Publishers, 2002.
[Har65]
Hartmanis, J., and Stearns, R.E. On the computational complexity of
algorithms. Transactions of the American Mathematical Society, vol.
117,1965,285-306.
[Har92]
Hare!, D. Algorithmics: The Spirit of Computing, 2nd ed. AddisonWesley, Reading, MA, 1992.
[Hea63]
Heap, B.R. Permutations by interchanges. Computer Journal, vol. 6,
1963, 293-294.
[Hoa96]
Hoare, C.A.R. Quicksort. In Great Papers in Computer Science,
Phillip Laplante, ed. West Publishing Company, Minneapolis/St.
Paul, 1996, pp. 31-39.
[Hoc97]
Hochbaum, D.S., editor. Approximation Algorithms for NP-Hard
Problems. PWS Publishing, Boston, 1997.
[Hop73]
Hopcroft, J.E. and Karp, R.M. A n 512 algorithm for maximum
matching in bipartite graphs. SIAM Journal on Computing, vol.
2, 1973, 225-231.
[Hop87]
Hopcroft, J.E. Computer science:· The emergence of a discipline.
Communications of the ACM, vol. 30, no. 3, 1987, 198--202.
[Hor80]
Horspool, R.N. Practical fast searching in strings. Software-Practice
and Experience, vol. 10, 1980,501-506.
[Hor98]
Horowitz, E., Sahni, S., and Rajasekaran, S. Computer Algorithms.
Computer Science Press, New York, 1998.
[Hu02]
Hu, T.C. and Shing, M.T. Combinatorial Algorithms. Enlarged 2nd
edition. Dover Publications, Mineola, NY, 2002.
[Huf52]
Huffman, D.A. A method for the construction of minimum redundancy codes. Proceedings of the IRE, vol. 40, 1952, 1098-1101.
[Joh02]
Johnson, D.S. and McGeoch, L.A. Experimental analysis of heuristics
for the STSP. In Traveling Salesman Problem and Its Variations,
edited by G. Gutin and A.P. Punnen, Kluwer Academic Publishers,
2002.
[Joh97]
Johnson, D.S. and McGeoch, L.A. The traveling salesman problem: a
case study. In Local Search in Combinatorial Optimization, edited by
II
492
Bibliography
[Kar72]
[Kar84]
[Kar86]
E. Aarts and J.K. Lenstra. John Wiley and Sons, Chichester, England,
1997.
Karp, R.M. Reducibility among combinatorial problems. In Complexity of Computer Communications, edited by R.E. Miller and J.W.
Thatcher. Plenum Press, New York, 1972,85-103.
.
Karmarkar, N. A new polynomial-time algorithm for linear programming. Combinatorica, vol. 4, no. 4, 1984, 373-395.
Karp, R.M. Combinatorics, complexity, and randomness. Communications of the ACM, vol. 29, no. 2, 1986, 89-109.
[Kar91]
Karp, R.M. An introduction to randomized algorithms. Discrete
Applied Mathematics, vol. 34, 1991, 165-201.
[Ker99]
Kernighan, B.W., and Pike. R. The Practice of Programming.
Addison-Wesley, Reading, MA, 1999.
Khachian, L.G. A polynomial algorithm in linear programming.
Soviet Mathematics Doklady, vol. 20, 1979, 191-194.
[Kha02]
[Knu75]
Knuth, D.E. Estimating the efficiency of backtrack programs.
Mathematics of Computation, vol. 29,1975,121-136.
[Knu76]
Knuth, D. E. Big omicron and big omega and big theta. ACM SIGACT
News, vol. 8, no. 2, 1976, 18-23.
[Knu77]
Knuth, D.E., Morris, Jr., J.H., and Pratt, V.R. Fast pattern matching
in strings. SIAM Journal on Computing, vol. 5, no. 2, 1977, 323-350.
[Knu96]
Knuth, D.E. Selected Papers on Computer Science. CSLI Publications
and Cambridge University Press, New York, 1996.
Knuth, D.E. The Art of Computer Programming, Volume 1: Fundamental Algorithms, 3rd ed. Addison-Wesley, Reading, MA, 1997.
[Knui]
[Knuii]
Knuth, D.E. The Art of Computer Programming, Volume 2: Seminumerical Algorithms, 3rd ed. Addison-Wesley, Reading, MA, 1998.
[Knuiii]
Knuth, D.E. The Art of Computer Programming, Volume 3: Sorting
and Searching, 2nd ed. Addison-Wesley, Reading, MA, 1998.
[Kol95]
Kolman, B., and Beck, R.E. Elementary Linear Programming with
Applications, 2nd ed. Academic Press, San Diego, 1995.
[Kor92]
Kordemsky, B.A. The Moscow Puzzles. Dover Publications, New
York, 1992.
Kruskal, J.B. On the shortest spanning subtree of a graph and
the traveling salesman problem. Proceedings of the American
Mathematical Society, vol. 7, 1956,48-50.
[Kru56]
[Law85]
Lawler, E.L., Lenstra, J.K., Rinnooy Kan, A.H.G., and Shmoys, D.B.,
editors. The Traveling Salesman Problem. John Wiley & Sons, New
York, 1985.
i
T!
'"<->/
l
I
I
Bibliography
493
(Lev73]
Levin, L.A. Universal sorting problems. Problemy Peredachi lnformatsii, val. 9, no. 3, 1973, 115-116 (in Russian). English translation in
Problems of Information Transmission, val. 9, 265-266.
(Lev99]
Levitin, A. Do we teach the right algorithm design techniques? In
Proceedings SIGCSE '99, New Orleans, 1999, 179-183.
[Lin73]
Lin, S. and Kernighan, B.W. An effective heuristic algorithm for
the traveling-salesman problem. Operations Research, val. 21, 1973,
498-516.
[Man89]
Manber, U. Introduction to Algorithms: A Creative Approach.
Addison-Wesley, Reading, MA, 1989.
[Mar90]
Martello, S., and Toth, P. Knapsack Problems: Algorithms and
Computer Implementations. John Wiley & Sons, Chichester, England,
1990.
[MicOO]
Michalewicz, Z. and Fogel, D.B. How to Solve It: Modern Heuristics.
Springer-Verlag, 2000.
[MilOO]
Miller, R., and Boxer, L. A Unified Approach to Sequential and
Parallel Algorithms. Prentice Hall, Upper Saddle River, NJ, 2000.
(Mor91]
Morel, B.M.E., and Shapiro, H.D. Algorithms from P to NP, Volume
I: Design and Efficiency. Benjamin Cummings, Redwood City, CA,
1991.
[MosOl]
Moscovich, I. 1000 Play Thinks: Puzzle.~ Paradoxes, Illusions, and
Games. Workman, New York, 2001.
[Mot95]
Motwani, R., and Raghavan, P. Randomized Algorithms. Cambridge
University Press, New York, 1995.
[Nea04]
Neapolitan, R., and Naimipour, K. Foundations of Algorithms: Using
Java Pseudocode. Jones and Bartlett, Sudbury, MA, 2004.
[Nem89]
Nemhauser, G.L., A.H.G. Rinnooy Kan, and M.J. Todd (eds.).
Optimization. North-Holland, Amsterdam, 1989.
[NieOO]
Nielsen, M.A., and Chuang, I.L. Quantum Computation and
Quantum Information. Cambridge University Press, Cambridge,
UK,2000.
[0Co98]
O'Connor, J.J., and Robertson, E.F. The Mac Tutor History of Mathematics Archive, June 1998, http://www-history.mcs.st-andrews.ac.uk/
history /Mathematicians/Abel.html.
[Ong84]
Ong, H.L. and Moore, J.B. Worst-case analysis of two traveling
salesman heuristics. Operations Research Letters, val. 2, 1984, 273277.
[0Ro98]
!
J
O'Rourke, J. Computational Geometry in C. 2nd ed. Cambridge
University Press, Cambridge, UK, 1998.
il
~
f :;:
..Ji-t.));,!
'.'. ·.i
•.
'f
494
Bibliography
[Ove80]
Overmars, M.H., and van Leeuwen, J. Further Comments on Bykat's
Convex Hull Algorithm. Inj(>rmation Processing Letters, vol. 10, no.
4/5, 1980, 209-212.
[Pan78]
Pan, V.Y. Strassen's algorithm is not optimal. In Proceedings of
Ninteenth Annual IEEE Symposium on the Foundations of Computer
Science.1978, 166-176.
[Pap82]
Papadimitriou, C.H., and Steiglitz, K. Combinatorial Optimization:
Algorithms and Complexity. Prentice-Hall, Englewood Cliffs, NJ,
1982.
[Pap94]
Papadimitriou, C. H. Computational Complexity. Addison-Wesley,
Reading, MA, 1994.
[Par95]
Parberry, I. Problems on Algorithms. Prentice-Hall, Englewood
Cliffs, NJ, 1995.
[Pol57]
Polya, G. How to Solve It. Doubleday, Garden City, NY, 1957.
[Pre85]
Preparata, F.P., and Shamos, M.l. Computational Geometry: An
Introduction. Springer-Verlag, New York, 1985.
[Pri57]
Prim, R.C. Shortest connection networks and some generalizations.
Bell System Technicallournal, vol. 36, no. 1, 1957, 1389-1401.
[Pri93]
Price B., Baecker, R., and Small, l. A principled taxonomy of software
visualization. Journal of Visual Languages and Computing, vol. 4, no.
3, 1993, 211-266.
[Pur85]
Purdom, P.W., Jr., and Brown, C. The Analysis of Algorithms. Holt,
Rinehart and Winston, New York, 1985.
[Raw91]
Rawlins, G.J.E. Compared to What? An Introduction to the Analysis
of Algorithms. Computer Science Press, New York, 1991.
[Rei77]
Reingold, E.M., Nievergelt, J., and Deo, N. Combinatorial Algorithms: Theory and Practice. Prentice-Hall, Englewood Cliffs, NJ,
1977.
[Riv78]
Rivest, R.L., Sbamir, A., and Adleman, L.M. A method for obtaining
digital signatures and public-key cryptosystems. Communications of
the ACM, vol. 21, no. 2, 1978, 120-126.
[Ros77]
Rosenkrantz, D.J., Stearns, R.E., and Lewis, P.M. An analysis of
several heuristics for the traveling salesman problem. SIAM Journal
of Computing, vol. 6,1977,563-581.
[Sah75]
Sahni, S. Approximation algorithms for the 0/1 knapsack problem.
Journal of the ACM, vol. 22,1975,115-124.
[Sah76]
Sahni, S., and Gonzalez, T. ?-complete approximation problems.
Journal of the ACM, vol. 23,1976,555-565.
[Sav03]
Savchev, S., and Andreescu, T. Mathematical miniatures. Mathematical Association of America, Washington, DC, 2003.
Bibliography
495
[SayOOJ
Sayood, K. Introduction to Data Compression, 2nd ed. Morgan
Kaufmann Publishers, San Francisco, CA, 2000.
[Sed88]
Sedgewick, R. Algorithms, 2nd ed. Addison-Wesley, Reading, MA,
1988.
[Sed96]
Sedgewick, R., and Flajolet, P An Introduction to the Analysis of
Algorithms. Addison-Wesley, Reading, MA, 1996.
[Sed98]
Sedge wick, R. Algorithms inc; 3rd ed. Parts 1-5: Fundamentals, Data
Structnres, Sorting, Searching, and Graph Algorithms. AddisonWesley, Reading, MA, 1998.
[Sha98]
Shasha, D., and Lazere, C. Out of Their Minds: The Lives and
Discoveries of 15 Great Computer Scientists. Copernicus, New York,
1998.
[She59]
Shell, D.L. A high-speed sorting procedure. Communications of the
ACM, vol. 2, no. 7, July 1959, 30-32.
[She97]
Shen, A. Algorithms and Programming: Problems and Solutions.
Birkhauscr, Boston, 1997.
[Sho94]
Shor, PW. Algorithms for quantum computation: Discrete algorithms
and factoring. Proceedings of the 35th Annual Symposium on
Foundations of Computer Science (Shafi Goldwasser, editor). IEEE
Computer Society Press, 1994, 124-134.
Sipser, M. Introduction to the Theory of Computation. PSW Publishing, Boston, 1997.
Skiena, S.S. Algorithm Design Manual. Springer-Verlag, New York,
1998.
Stasko, J. TANGO: A framework and system for algorithm animation. Computer, vol. 23, no. 9, 1990, 27-39.
Strassen, V. Gaussian elimination is not optimal. Numerische
Mathematik, vol. 13, 1969,354-356.
Tarjan, R.E. Data Structures and Network Algorithms. Society for
Industrial and Applied Mathematics, Philadelphia, 1983.
Tarjan, R.E., and van Leeuwen, J. Worst-case analysis of set union
algorithms. Journal of the ACM, vol. 31, no. 2, 1984, 245-281.
Tarjan, R.E. Amortized computational complexity. SIAM Journal on
Algebraic and Discrete Methods, vol. 6, no. 2, 1985, 306-318.
Tarjan, R.E. Algorithm design. Communications of the ACM, vol.
30, no. 3, 1987, 204-212.
Tucker, A. Applied Combinatorics. John Wiley & Sons, New York,
1980.
:
I:
r
ji
496
Bibliography
[War62]
Warshall, S. A theorem on boolean matrices. Journal of the ACM,
vol. 9, no. 1, 1962, 11-12.
[Wei77]
Weide, B. A survey of analysis techniques for discrete algorithms.
Computing Surveys, vol. 9, no. 4, 1977,291-313.
[Wei98]
Weiss, M.A. Data Structures and Algorithm Analysis in Java,
Addison-Wesley, 1998.
[Wil64]
Williams, J.W.J. Algorithm 232 (heapsort). Communications of the
ACM, vol. 7, no. 6, 1964, 347-348.
Wirth, N. Algorithms + Data Structures= Programs. Prentice-Hall,
Englewood Cliffs, NJ, 1976.
Yao, F. Speed-up in dynamic programming. SIAM Journal on
Algebraic and Discrete Methods, vol. 3, no. 4, 1982, 532-540.
Ye, Yinyu. Interior Point Algorithms: Theory and Analysis. John
Wiley & Sons, 1997.
[Wir76]
[Yao82]
[Ye97]
Hints to Exercises
CHAPTER 1
Exercises 1.1
1. It is probably faster to do this by searching the Web, but your library should
be able to help, too.
2. Oue can find arguments supporting either view. There is a well-established
principle pertinent to the matter, though: scientific facts or mathematical
expressions of them are not patentable. (Why do you think this is the case?)
But should this preclude granting patents for all algorithms?
3. You may assume that you are writing your algorithms for a human rather
than a machine. Still, make sure that your descriptions do not contain obvious
ambiguities. Knuth ([Knui], p. 6) provides an interesting comparison between
cooking recipes and algorithms.
4. There is a quite straightforward algorithm for this problem based on the
definition of L,fiij.
5. a. Just follow Euclid's algorithm as described in the text.
b. Compare the number of divisions made by the two algorithms.
6. Prove that if d divides both m and n (i.e., m = sd and n = td for some positive
integers s and t), then it also divides both n and r = m mod n and vice versa.
Use the formula m = qn + r (0 :S r < n) and the fact that if d divides two
integers u and v, it also divides u + v and u ~ v. (Why?)
7. Perform one iteration of the algorithm for two arbitrarily chosen integers
m <n.
8. The answer to part (a) can be given immediately; the answer to part (b) can
be given by checking the algorithm's performance on all pairs 1 < m < n :S 10.
9. a. Use the equality
gcd(m, n) = gcd(m ~ n, n) form C': n > 0.
497
498
Hints to Exercises
b. The key is to figure out the total number of distinct numbers that can be
written on the board, starting with an initial pair m, n where m > n ::0: L
You should exploit a connection of this question to the question of part (a).
Consideri